Using the transaction-begin instruction to manage transactional aborts in transactional memory computing environments

ABSTRACT

When executed, a transaction-begin instruction specifies an initial value for a transaction-count-to-completion (CTC) value for a transaction. The initial value indicates a predicted duration of the transaction. The CTC value may be a number of instructions to completion or an amount of time to completion. The CTC value is adjusted as the transaction progresses. The adjusted CTC value indicates how far the transaction is from completion. When a disruptive event associated with inducing transactional aborts, such as an interrupt or a conflicting memory access, is identified while processing the transaction, processing of the disruptive event is deferred if the adjusted CTC value satisfies deferral criteria. If the adjusted CTC value does not satisfy deferral criteria, the transaction is aborted and the disruptive event is processed.

BACKGROUND

This disclosure relates generally to transaction processing in amulti-processor computing environment with transactional memory, andmore specifically to using the transaction-begin instruction to managetransactional aborts in a transactional memory computing environment.

The number of central processing unit (CPU) cores on a chip and thenumber of CPU cores connected to a shared memory continue to growsignificantly to support growing workload capacity demand. Theincreasing number of CPUs cooperating to process the same workloads putsa significant burden on software scalability; for example, shared queuesor data-structures protected by traditional semaphores become hot spotsand lead to sub-linear n-way scaling curves. Traditionally this has beencountered by implementing finer-grained locking in software, and withlower latency/higher bandwidth interconnects in hardware. Implementingfine-grained locking to improve software scalability can be verycomplicated and error-prone, and at today's CPU frequencies, thelatencies of hardware interconnects are limited by the physicaldimension of the chips and systems, and by the speed of light.

Implementations of hardware Transactional Memory (HTM, or in thisdiscussion, simply TM) have been introduced, wherein a group ofinstructions—called a transaction—operate in an atomic manner on a datastructure in memory, as viewed by other central processing units (CPUs)and the I/O subsystem (atomic operation is also known as “blockconcurrent” or “serialized” in other literature). The transactionexecutes optimistically without obtaining a lock, but may need to abortand retry the transaction execution if an operation of the executingtransaction on a memory location conflicts with another operation on thesame memory location. Previously, software transactional memoryimplementations have been proposed to support software TransactionalMemory (TM). However, hardware TM can provide improved performanceaspects and ease of use over software TM.

U.S. Patent Application Publication No. 2011/0145498 titled“Instrumentation of Hardware Assisted Transactional Memory System” filed2009 Dec. 15, incorporated herein by reference in its entirety, teachesmonitoring performance of one or more architecturally significantprocessor caches coupled to a processor. The methods include executingan application on one or more processors coupled to one or morearchitecturally significant processor caches, where the applicationutilizes the architecturally significant portions of the architecturallysignificant processor caches. The methods further include at least oneof generating metrics related to performance of the architecturallysignificant processor caches; implementing one or more debug exceptionsrelated to performance of the architecturally significant processorcaches; or implementing one or more transactional breakpoints related toperformance of the architecturally significant processor caches as aresult of utilizing the architecturally significant portions of thearchitecturally significant processor caches.

U.S. Patent Application Publication No. 2011/0119452 titled “HybridTransactional Memory System (Hybrid TM) and Method” filed 2009 Nov. 16,incorporated herein by reference in its entirety, teaches a computerprocessing system having memory and processing facilities for processingdata with a computer program that is a Hybrid Transactional Memorymultiprocessor system with modules 1 . . . n coupled to a systemphysical memory array, I/O devices via a high speed interconnectionelement. A CPU is integrated as in a multi-chip module withmicroprocessors which contain or are coupled in the CPU module to anassist thread facility, as well as a memory controller, cachecontrollers, cache memory, and other components which form part of theCPU which connects to the high speed interconnect which functions underthe architecture and operating system to interconnect elements of thecomputer system with physical memory, various I/O devices, and the otherCPUs of the system. The hybrid transactional memory elements support fora transactional memory system that has a simple/cost effective hardwaredesign that can deal with limited hardware resources, yet one which hasa transactional facility control logic providing for a back up assistthread that can still allow transactions to reference existing librariesand allows programmers to include calls to existing software librariesinside of their transactions, and which will not make a user code use asecond lock based solution.

SUMMARY

Disclosed herein are embodiments of a method for processing transactionsin a transaction execution computing environment with transactionalmemory. A transaction-begin instruction associated with a transaction isexecuted in the computing environment. The transaction-begin instructionspecifies an initial value for a transaction-count-to-completion (CTC)value. The initial value indicates a predicted duration of thetransaction.

The CTC value is adjusted based on the progress of the transaction. Theadjusted CTC value predicts how far the transaction is from completion.In some embodiments, the CTC value indicates the predicted number ofinstructions remaining to complete execution of the transaction. In someembodiments, the CTC value indicates the predicted amount of timeremaining to complete execution of the transaction. In some embodiments,the CTC value may indicate either the predicted number of instructionsor the predicted amount of time remaining to complete the transaction.

When a pending disruptive event associated with inducing transactionalaborts is identified while processing the transaction, such as a pendingasynchronous interrupt or a conflicting memory access of memorymonitored by the transaction, processing of the identified pendingdisruptive event is deferred based on determining that the adjusted CTCvalue satisfies predefined deferral criteria.

Also disclosed herein are embodiments of a computer system with a memoryand a processor in communication with the memory. The computer system isconfigured to perform the method described above. Also disclosed hereinare embodiments of a computer program product with a computer readablestorage medium readable by a processing circuit and storing instructionsfor execution by the processing circuit. The instructions are forperforming the method described above.

BRIEF DESCRIPTION OF THE SEVERAL VIEWS OF THE DRAWINGS

One or more aspects of the present disclosed embodiments areparticularly pointed out and distinctly claimed as examples in theclaims at the conclusion of the specification. The foregoing and otherobjects, features, and advantages of the disclosed embodiments areapparent from the following detailed description taken in conjunctionwith the accompanying drawings in which:

FIG. 1 depicts an example multicore Transactional Memory environment, inaccordance with embodiments of the present disclosure;

FIG. 2 depicts example components of an example CPU, in accordance withembodiments of the present disclosure;

FIG. 3 is a block diagram representing an example system for managingtransactional aborts with a transaction-begin instruction, in accordancewith embodiments of the present disclosure;

FIG. 4 is a flow diagram illustrating an example method for managingtransactional aborts with a transaction-begin instruction, in accordancewith embodiments of the present disclosure;

FIG. 5 is a flow diagram illustrating an example method for processing atime-to-completion transaction-begin instruction for a transaction, inaccordance with embodiments of the present disclosure;

FIG. 6 is a flow diagram illustrating an example method for processingan instructions-to-completion transaction-begin instruction for atransaction, in accordance with embodiments of the present disclosure;

FIG. 7 is a block diagram illustrating an example system for managingtransactional aborts with a transaction-begin instruction, in accordancewith embodiments of the present disclosure;

FIG. 8 is a block diagram of an example emulated host computer systemfor managing transactional aborts with a transaction-begin instruction,in accordance with embodiments of the present disclosure; and

FIG. 9 is a block diagram of an example computer program product formanaging transactional aborts with a transaction-begin instruction, inaccordance with embodiments of the present disclosure.

DETAILED DESCRIPTION

Historically, a computer system or processor had only a single processor(also known as processing unit or central processing unit). Theprocessor included an instruction processing unit (IPU), a branch unit,a memory control unit and the like. Such processors were capable ofexecuting a single thread of a program at a time. Operating systems weredeveloped that could time-share a processor by dispatching a program tobe executed on the processor for a period of time, and then dispatchinganother program to be executed on the processor for another period oftime. As technology evolved, memory subsystem caches were often added tothe processor as well as complex dynamic address translation includingtranslation lookaside buffers (TLBs). The IPU itself was often referredto as a processor. As technology continued to evolve, an entireprocessor could be packaged in a single semiconductor chip or die; sucha processor was referred to as a microprocessor. Then processors weredeveloped that incorporated multiple IPUs; such processors were oftenreferred to as multi-processors. Each such processor of amulti-processor computer system (processor) may include individual orshared caches, memory interfaces, system bus, address translationmechanism and the like. Virtual machine and instruction set architecture(ISA) emulators added a layer of software to a processor that providedthe virtual machine with multiple “virtual processors” (also known asprocessors) by time-slice usage of a single IPU in a single hardwareprocessor. As technology further evolved, multi-threaded processors weredeveloped, enabling a single hardware processor having a singlemulti-thread IPU to provide a capability of simultaneously executingthreads of different programs, thus each thread of a multi-threadedprocessor appeared to the operating system as a processor. As technologyfurther evolved, it was possible to put multiple processors (each havingan IPU) on a single semiconductor chip or die. These processors werereferred to processor cores or just cores. Thus the terms such asprocessor, central processing unit, processing unit, microprocessor,core, processor core, processor thread, and thread, for example, areoften used interchangeably. Aspects of embodiments herein may bepracticed by any or all processors including those shown supra, withoutdeparting from the teachings herein. Wherein the term “thread” or“processor thread” is used herein, it is expected that particularadvantage of the embodiment may be had in a processor threadimplementation.

Transaction Execution in Intel® Based Embodiments

In “Intel® Architecture Instruction Set Extensions ProgrammingReference” 319433-012A, February 2012, incorporated herein by referencein its entirety, Chapter 8 teaches, in part, that multithreadedapplications may take advantage of increasing numbers of CPU cores toachieve higher performance. However, the writing of multi-threadedapplications requires programmers to understand and take into accountdata sharing among the multiple threads. Access to shared data typicallyrequires synchronization mechanisms. These synchronization mechanismsare used to ensure that multiple threads update shared data byserializing operations that are applied to the shared data, oftenthrough the use of a critical section that is protected by a lock. Sinceserialization limits concurrency, programmers try to limit the overheaddue to synchronization.

Intel® Transactional Synchronization Extensions (Intel® TSX) allow aprocessor to dynamically determine whether threads need to be serializedthrough lock-protected critical sections, and to perform thatserialization only when required. This allows the processor to exposeand exploit concurrency that is hidden in an application because ofdynamically unnecessary synchronization.

With Intel TSX, programmer-specified code regions (also referred to as“transactional regions” or just “transactions”) are executedtransactionally. If the transactional execution completes successfully,then all memory operations performed within the transactional regionwill appear to have occurred instantaneously when viewed from otherprocessors. A processor makes the memory operations of the executedtransaction performed within the transactional region visible to otherprocessors only when a successful commit occurs, i.e., when thetransaction successfully completes execution. This process is oftenreferred to as an atomic commit.

Intel TSX provides two software interfaces to specify regions of codefor transactional execution. Hardware Lock Elision (HLE) is a legacycompatible instruction set extension (comprising the XACQUIRE andXRELEASE prefixes) to specify transactional regions. RestrictedTransactional Memory (RTM) is a new instruction set interface(comprising the XBEGIN, XEND, and XABORT instructions) for programmersto define transactional regions in a more flexible manner than thatpossible with HLE. HLE is for programmers who prefer the backwardcompatibility of the conventional mutual exclusion programming model andwould like to run HLE-enabled software on legacy hardware but would alsolike to take advantage of the new lock elision capabilities on hardwarewith HLE support. RTM is for programmers who prefer a flexible interfaceto the transactional execution hardware. In addition, Intel TSX alsoprovides an XTEST instruction. This instruction allows software to querywhether the logical processor is transactionally executing in atransactional region identified by either HLE or RTM.

Since a successful transactional execution ensures an atomic commit, theprocessor executes the code region optimistically without explicitsynchronization. If synchronization was unnecessary for that specificexecution, execution can commit without any cross-thread serialization.If the processor cannot commit atomically, then the optimistic executionfails. When this happens, the processor will roll back the execution, aprocess referred to as a transactional abort. On a transactional abort,the processor will discard all updates performed in the memory regionused by the transaction, restore architectural state to appear as if theoptimistic execution never occurred, and resume executionnon-transactionally.

A processor can perform a transactional abort for numerous reasons. Aprimary reason to abort a transaction is due to conflicting memoryaccesses between the transactionally executing logical processor andanother logical processor. Such conflicting memory accesses may preventa successful transactional execution. Memory addresses read from withina transactional region constitute the read-set of the transactionalregion and addresses written to within the transactional regionconstitute the write-set of the transactional region. Intel TSXmaintains the read- and write-sets at the granularity of a cache line. Aconflicting memory access occurs if another logical processor eitherreads a location that is part of the transactional region's write-set orwrites a location that is a part of either the read- or write-set of thetransactional region. A conflicting access typically means thatserialization is required for this code region. Since Intel TSX detectsdata conflicts at the granularity of a cache line, unrelated datalocations placed in the same cache line will be detected as conflictsthat result in transactional aborts. Transactional aborts may also occurdue to limited transactional resources. For example, the amount of dataaccessed in the region may exceed an implementation-specific capacity.Additionally, some instructions and system events may causetransactional aborts. Frequent transactional aborts result in wastedcycles and increased inefficiency.

Hardware Lock Elision

Hardware Lock Elision (HLE) provides a legacy compatible instruction setinterface for programmers to use transactional execution. HLE providestwo new instruction prefix hints: XACQUIRE and XRELEASE.

With HLE, a programmer adds the XACQUIRE prefix to the front of theinstruction that is used to acquire the lock that is protecting thecritical section. The processor treats the prefix as a hint to elide thewrite associated with the lock acquire operation. Even though the lockacquire has an associated write operation to the lock, the processordoes not add the address of the lock to the transactional region'swrite-set nor does it issue any write requests to the lock. Instead, theaddress of the lock is added to the read-set. The logical processorenters transactional execution. If the lock was available before theXACQUIRE prefixed instruction, then all other processors will continueto see the lock as available afterwards. Since the transactionallyexecuting logical processor neither added the address of the lock to itswrite-set nor performed externally visible write operations to the lock,other logical processors can read the lock without causing a dataconflict. This allows other logical processors to also enter andconcurrently execute the critical section protected by the lock. Theprocessor automatically detects any data conflicts that occur during thetransactional execution and will perform a transactional abort ifnecessary.

Even though the eliding processor did not perform any external writeoperations to the lock, the hardware ensures program order of operationson the lock. If the eliding processor itself reads the value of the lockin the critical section, it will appear as if the processor had acquiredthe lock, i.e. the read will return the non-elided value. This behaviorallows an HLE execution to be functionally equivalent to an executionwithout the HLE prefixes.

An XRELEASE prefix can be added in front of an instruction that is usedto release the lock protecting a critical section. Releasing the lockinvolves a write to the lock. If the instruction is to restore the valueof the lock to the value the lock had prior to the XACQUIRE prefixedlock acquire operation on the same lock, then the processor elides theexternal write request associated with the release of the lock and doesnot add the address of the lock to the write-set. The processor thenattempts to commit the transactional execution.

With HLE, if multiple threads execute critical sections protected by thesame lock but they do not perform any conflicting operations on eachother's data, then the threads can execute concurrently and withoutserialization. Even though the software uses lock acquisition operationson a common lock, the hardware recognizes this, elides the lock, andexecutes the critical sections on the two threads without requiring anycommunication through the lock—if such communication was dynamicallyunnecessary.

If the processor is unable to execute the region transactionally, thenthe processor will execute the region non-transactionally and withoutelision. HLE enabled software has the same forward progress guaranteesas the underlying non-HLE lock-based execution. For successful HLEexecution, the lock and the critical section code must follow certainguidelines. These guidelines only affect performance; failure to followthese guidelines will not result in a functional failure. Hardwarewithout HLE support will ignore the XACQUIRE and XRELEASE prefix hintsand will not perform any elision since these prefixes correspond to theREPNE/REPE IA-32 prefixes which are ignored on the instructions whereXACQUIRE and XRELEASE are valid. Importantly, HLE is compatible with theexisting lock-based programming model. Improper use of hints will notcause functional bugs though it may expose latent bugs already in thecode.

Restricted Transactional Memory (RTM) provides a flexible softwareinterface for transactional execution. RTM provides three newinstructions—XBEGIN, XEND, and XABORT—for programmers to start, commit,and abort a transactional execution.

The programmer uses the XBEGIN instruction to specify the start of atransactional code region and the XEND instruction to specify the end ofthe transactional code region. If the RTM region could not besuccessfully executed transactionally, then the XBEGIN instruction takesan operand that provides a relative offset to the fallback instructionaddress.

A processor may abort RTM transactional execution for many reasons. Inmany instances, the hardware automatically detects transactional abortconditions and restarts execution from the fallback instruction addresswith the architectural state corresponding to that present at the startof the XBEGIN instruction and the EAX register updated to describe theabort status.

The XABORT instruction allows programmers to abort the execution of anRTM region explicitly. The XABORT instruction takes an 8-bit immediateargument that is loaded into the EAX register and will thus be availableto software following an RTM abort. RTM instructions do not have anydata memory location associated with them. While the hardware providesno guarantees as to whether an RTM region will ever successfully committransactionally, most transactions that follow the recommendedguidelines are expected to successfully commit transactionally. However,programmers must always provide an alternative code sequence in thefallback path to guarantee forward progress. This may be as simple asacquiring a lock and executing the specified code regionnon-transactionally. Further, a transaction that always aborts on agiven implementation may complete transactionally on a futureimplementation. Therefore, programmers must ensure the code paths forthe transactional region and the alternative code sequence arefunctionally tested.

Detection of HLE Support

A processor supports HLE execution if CPUID.07H.EBX.HLE [bit 4]=1.However, an application can use the HLE prefixes (XACQUIRE and XRELEASE)without checking whether the processor supports HLE. Processors withoutHLE support ignore these prefixes and will execute the code withoutentering transactional execution.

Detection of RTM Support

A processor supports RTM execution if CPUID.07H.EBX.RTM [bit 11]=1. Anapplication must check if the processor supports RTM before it uses theRTM instructions (XBEGIN, XEND, XABORT). These instructions willgenerate a #UD exception when used on a processor that does not supportRTM.

Detection of XTEST Instruction

A processor supports the XTEST instruction if it supports either HLE orRTM. An application must check either of these feature flags beforeusing the XTEST instruction. This instruction will generate a #UDexception when used on a processor that does not support either HLE orRTM.

Querying Transactional Execution Status

The XTEST instruction can be used to determine the transactional statusof a transactional region specified by HLE or RTM. Note, while the HLEprefixes are ignored on processors that do not support HLE, the XTESTinstruction will generate a #UD exception when used on processors thatdo not support either HLE or RTM.

Requirements for HLE Locks

For HLE execution to successfully commit transactionally, the lock mustsatisfy certain properties and access to the lock must follow certainguidelines.

An XRELEASE prefixed instruction must restore the value of the elidedlock to the value it had before the lock acquisition. This allowshardware to safely elide locks by not adding them to the write-set. Thedata size and data address of the lock release (XRELEASE prefixed)instruction must match that of the lock acquire (XACQUIRE prefixed) andthe lock must not cross a cache line boundary.

Software should not write to the elided lock inside a transactional HLEregion with any instruction other than an XRELEASE prefixed instruction,otherwise such a write may cause a transactional abort. In addition,recursive locks (where a thread acquires the same lock multiple timeswithout first releasing the lock) may also cause a transactional abort.Note that software can observe the result of the elided lock acquireinside the critical section. Such a read operation will return the valueof the write to the lock.

The processor automatically detects violations to these guidelines, andsafely transitions to a non-transactional execution without elision.Since Intel TSX detects conflicts at the granularity of a cache line,writes to data collocated on the same cache line as the elided lock maybe detected as data conflicts by other logical processors eliding thesame lock.

Transactional Nesting

Both HLE and RTM support nested transactional regions. However, atransactional abort restores state to the operation that startedtransactional execution: either the outermost) (ACQUIRE prefixed HLEeligible instruction or the outermost XBEGIN instruction. The processortreats all nested transactions as one transaction.

HLE Nesting and Elision

Programmers can nest HLE regions up to an implementation specific depthof MAX_HLE_NEST_COUNT. Each logical processor tracks the nesting countinternally but this count is not available to software. An XACQUIREprefixed HLE-eligible instruction increments the nesting count, and anXRELEASE prefixed HLE-eligible instruction decrements it. The logicalprocessor enters transactional execution when the nesting count goesfrom zero to one. The logical processor attempts to commit only when thenesting count becomes zero. A transactional abort may occur if thenesting count exceeds MAX_HLE_NEST_COUNT.

In addition to supporting nested HLE regions, the processor can alsoelide multiple nested locks. The processor tracks a lock for elisionbeginning with the XACQUIRE prefixed HLE eligible instruction for thatlock and ending with the XRELEASE prefixed HLE eligible instruction forthat same lock. The processor can, at any one time, track up to aMAX_HLE_ELIDED_LOCKS number of locks. For example, if the implementationsupports a MAX_HLE_ELIDED_LOCKS value of two and if the programmer neststhree HLE identified critical sections (by performing XACQUIRE prefixedHLE eligible instructions on three distinct locks without performing anintervening XRELEASE prefixed HLE eligible instruction on any one of thelocks), then the first two locks will be elided, but the third won't beelided (but will be added to the transaction's writeset). However, theexecution will still continue transactionally. Once an XRELEASE for oneof the two elided locks is encountered, a subsequent lock acquiredthrough the XACQUIRE prefixed HLE eligible instruction will be elided.

The processor attempts to commit the HLE execution when all elidedXACQUIRE and XRELEASE pairs have been matched, the nesting count goes tozero, and the locks have satisfied requirements. If execution cannotcommit atomically, then execution transitions to a non-transactionalexecution without elision as if the first instruction did not have anXACQUIRE prefix.

RTM Nesting

Programmers can nest RTM regions up to an implementation specificMAX_RTM_NEST_COUNT. The logical processor tracks the nesting countinternally but this count is not available to software. An XBEGINinstruction increments the nesting count, and an XEND instructiondecrements the nesting count. The logical processor attempts to commitonly if the nesting count becomes zero. A transactional abort occurs ifthe nesting count exceeds MAX_RTM_NEST_COUNT.

Nesting HLE and RTM

HLE and RTM provide two alternative software interfaces to a commontransactional execution capability. Transactional processing behavior isimplementation specific when HLE and RTM are nested together, e.g., HLEis inside RTM or RTM is inside HLE. However, in all cases, theimplementation will maintain HLE and RTM semantics. An implementationmay choose to ignore HLE hints when used inside RTM regions, and maycause a transactional abort when RTM instructions are used inside HLEregions. In the latter case, the transition from transactional tonon-transactional execution occurs seamlessly since the processor willre-execute the HLE region without actually doing elision, and thenexecute the RTM instructions.

Abort Status Definition

RTM uses the EAX register to communicate abort status to software.Following an RTM abort the EAX register has the following definition.

TABLE 1 RTM Abort Status Definition EAX Register Bit Position Meaning 0Set if abort caused by XABORT instruction 1 If set, the transaction maysucceed on retry, this bit is always clear if bit 0 is set 2 Set ifanother logical processor conflicted with a memory address that was partof the transaction that aborted 3 Set if an internal buffer overflowed 4Set if a debug breakpoint was hit 5 Set if an abort occurred duringexecution of a nested transaction 23:6 Reserved 31:24 XABORT argument(only valid if bit 0 set, otherwise reserved)

The EAX abort status for RTM only provides causes for aborts. It doesnot by itself encode whether an abort or commit occurred for the RTMregion. The value of EAX can be 0 following an RTM abort. For example, aCPUID instruction when used inside an RTM region causes a transactionalabort and may not satisfy the requirements for setting any of the EAXbits. This may result in an EAX value of 0.

RTM Memory Ordering

A successful RTM commit causes all memory operations in the RTM regionto appear to execute atomically. A successfully committed RTM regionconsisting of an XBEGIN followed by an XEND, even with no memoryoperations in the RTM region, has the same ordering semantics as a LOCKprefixed instruction.

The XBEGIN instruction does not have fencing semantics. However, if anRTM execution aborts, then all memory updates from within the RTM regionare discarded and are not made visible to any other logical processor.

RTM-Enabled Debugger Support

By default, any debug exception inside an RTM region will cause atransactional abort and will redirect control flow to the fallbackinstruction address with architectural state recovered and bit 4 in EAXset. However, to allow software debuggers to intercept execution ondebug exceptions, the RTM architecture provides additional capability.

If bit 11 of DR7 and bit 15 of the IA32_DEBUGCTL_MSR are both 1, any RTMabort due to a debug exception (#DB) or breakpoint exception (#BP)causes execution to roll back and restart from the XBEGIN instructioninstead of the fallback address. In this scenario, the EAX register willalso be restored back to the point of the XBEGIN instruction.

Programming Considerations

Typical programmer-identified regions are expected to transactionallyexecute and commit successfully. However, Intel TSX does not provide anysuch guarantee. A transactional execution may abort for many reasons. Totake full advantage of the transactional capabilities, programmersshould follow certain guidelines to increase the probability of theirtransactional execution committing successfully.

This section discusses various events that may cause transactionalaborts. The architecture ensures that updates performed within atransaction that subsequently aborts execution will never becomevisible. Only committed transactional executions initiate an update tothe architectural state. Transactional aborts never cause functionalfailures and only affect performance.

Instruction Based Considerations

Programmers can use any instruction safely inside a transaction (HLE orRTM) and can use transactions at any privilege level. However, someinstructions will always abort the transactional execution and causeexecution to seamlessly and safely transition to a non-transactionalpath.

Intel TSX allows for most common instructions to be used insidetransactions without causing aborts. The following operations inside atransaction do not typically cause an abort:

-   -   Operations on the instruction pointer register, general purpose        registers (GPRs) and the status flags (CF, OF, SF, PF, AF, and        ZF); and    -   Operations on XMM and YMM registers and the MXCSR register.

However, programmers must be careful when intermixing SSE and AVXoperations inside a transactional region. Intermixing SSE instructionsaccessing XMM registers and AVX instructions accessing YMM registers maycause transactions to abort. Programmers may use REP/REPNE prefixedstring operations inside transactions. However, long strings may causeaborts. Further, the use of CLD and STD instructions may cause aborts ifthey change the value of the DF flag. However, if DF is 1, the STDinstruction will not cause an abort. Similarly, if DF is 0, then the CLDinstruction will not cause an abort.

Instructions not enumerated here as causing aborts when used inside atransaction will typically not cause a transaction to abort (examplesinclude but are not limited to MFENCE, LFENCE, SFENCE, RDTSC, RDTSCP,etc.).

The following instructions will abort transactional execution on anyimplementation:

-   -   XABORT    -   CPUID    -   PAUSE

In addition, in some implementations, the following instructions mayalways cause transactional aborts. These instructions are not expectedto be commonly used inside typical transactional regions. However,programmers must not rely on these instructions to force a transactionalabort, since whether they cause transactional aborts is implementationdependent.

-   -   Operations on X87 and MMX architecture state. This includes all        MMX and X87 instructions, including the FXRSTOR and FXSAVE        instructions.    -   Update to non-status portion of EFLAGS: CLI, STI, POPFD, POPFQ,        CLTS.    -   Instructions that update segment registers, debug registers        and/or control registers: MOV to DS/ES/FS/GS/SS, POP        DS/ES/FS/GS/SS, LDS, LES, LFS, LGS, LSS, SWAPGS, WRFSBASE,        WRGSBASE, LGDT, SGDT, LIDT, SIDT, LLDT, SLDT, LTR, STR, Far        CALL, Far JMP, Far RET, IRET, MOV to DRx, MOV to        CR0/CR2/CR3/CR4/CR8 and LMSW.    -   Ring transitions: SYSENTER, SYSCALL, SYSEXIT, and SYSRET.    -   TLB and Cacheability control: CLFLUSH, INVD, WBINVD, INVLPG,        INVPCID, and memory instructions with a non-temporal hint        (MOVNTDQA, MOVNTDQ, MOVNTI, MOVNTPD, MOVNTPS, and MOVNTQ).    -   Processor state save: XSAVE, XSAVEOPT, and XRSTOR.    -   Interrupts: INTn, INTO.    -   IO: IN, INS, REP INS, OUT, OUTS, REP OUTS and their variants.    -   VMX: VMPTRLD, VMPTRST, VMCLEAR, VMREAD, VMWRITE, VMCALL,        VMLAUNCH, VMRESUME, VMXOFF, VMXON, INVEPT, and INVVPID.    -   SMX: GETSEC.    -   UD2, RSM, RDMSR, WRMSR, HLT, MONITOR, MWAIT, XSETBV, VZEROUPPER,        MASKMOVQ, and V/MASKMOVDQU.

Runtime Considerations

In addition to the instruction-based considerations, runtime events maycause transactional execution to abort. These may be due to data accesspatterns or micro-architectural implementation features. The followinglist is not a comprehensive discussion of all abort causes.

Any fault or trap in a transaction that must be exposed to software willbe suppressed. Transactional execution will abort and execution willtransition to a non-transactional execution, as if the fault or trap hadnever occurred. If an exception is not masked, then that un-maskedexception will result in a transactional abort and the state will appearas if the exception had never occurred.

Synchronous exception events (#DE, #OF, #NP, #SS, #GP, #BR, #UD, #AC,#XF, #PF, #NM, #TS, #MF, #DB, #BP/INT3) that occur during transactionalexecution may cause an execution not to commit transactionally, andrequire a non-transactional execution. These events are suppressed as ifthey had never occurred. With HLE, since the non-transactional code pathis identical to the transactional code path, these events will typicallyre-appear when the instruction that caused the exception is re-executednon-transactionally, causing the associated synchronous events to bedelivered appropriately in the non-transactional execution. Asynchronousevents (NMI, SMI, INTR, IPI, PMI, etc.) occurring during transactionalexecution may cause the transactional execution to abort and transitionto a non-transactional execution. The asynchronous events will be pendedand handled after the transactional abort is processed.

Transactions only support write-back cacheable memory type operations. Atransaction may always abort if the transaction includes operations onany other memory type. This includes instruction fetches to UC memorytype.

Memory accesses within a transactional region may require the processorto set the Accessed and Dirty flags of the referenced page table entry.The behavior of how the processor handles this isimplementation-specific. Some implementations may allow the updates tothese flags to become externally visible even if the transactionalregion subsequently aborts. Some Intel TSX implementations may choose toabort the transactional execution if these flags need to be updated.Further, a processor's page-table walk may generate accesses to its owntransactionally written but uncommitted state. Some Intel TSXimplementations may choose to abort the execution of a transactionalregion in such situations. Regardless, the architecture ensures that, ifthe transactional region aborts, then the transactionally written statewill not be made architecturally visible through the behavior ofstructures such as TLBs.

Executing self-modifying code transactionally may also causetransactional aborts. Programmers must continue to follow the Intelrecommended guidelines for writing self-modifying and cross-modifyingcode even when employing HLE and RTM. While an implementation of RTM andHLE will typically provide sufficient resources for executing commontransactional regions, implementation constraints and excessive sizesfor transactional regions may cause a transactional execution to abortand transition to a non-transactional execution. The architectureprovides no guarantee of the amount of resources available to dotransactional execution and does not guarantee that a transactionalexecution will ever succeed.

Conflicting requests to a cache line accessed within a transactionalregion may prevent the transaction from executing successfully. Forexample, if logical processor P0 reads line A in a transactional regionand another logical processor P1 writes line A (either inside or outsidea transactional region) then logical processor P0 may abort if logicalprocessor P1's write interferes with processor P0's ability to executetransactionally.

Similarly, if P0 writes line A in a transactional region and P1 reads orwrites line A (either inside or outside a transactional region), then P0may abort if P1's access to line A interferes with P0's ability toexecute transactionally. In addition, other coherence traffic may attimes appear as conflicting requests and may cause aborts. While thesefalse conflicts may happen, they are expected to be uncommon. Theconflict resolution policy to determine whether P0 or P1 aborts in theabove scenarios is implementation specific.

Generic Transaction Execution Embodiments:

According to “ARCHITECTURES FOR TRANSACTIONAL MEMORY”, a dissertationsubmitted to the Department of Computer Science and the Committee onGraduate Studies of Stanford University in partial fulfillment of therequirements for the Degree of Doctor of Philosophy, by Austen McDonald,June 2009, incorporated by reference herein in its entirety,fundamentally, there are three mechanisms needed to implement an atomicand isolated transactional region: versioning, conflict detection, andcontention management.

To make a transactional code region appear atomic, all the modificationsperformed by that transactional code region must be stored and keptisolated from other transactions until commit time. The system does thisby implementing a versioning policy. Two versioning paradigms exist:eager and lazy. An eager versioning system stores newly generatedtransactional values in place and stores previous memory values on theside, in what is called an undo-log. A lazy versioning system stores newvalues temporarily in what is called a write buffer, copying them tomemory only on commit. In either system, the cache is used to optimizestorage of new versions.

To ensure that transactions appear to be performed atomically, conflictsmust be detected and resolved. The two systems, i.e., the eager and lazyversioning systems, detect conflicts by implementing a conflictdetection policy, either optimistic or pessimistic. An optimistic systemexecutes transactions in parallel, checking for conflicts only when atransaction commits. A pessimistic system checks for conflicts at eachload and store. Similar to versioning, conflict detection also uses thecache, marking each line as either part of the read-set, part of thewrite-set, or both. The two systems resolve conflicts by implementing acontention management policy. Many contention management policies exist,some are more appropriate for optimistic conflict detection and some aremore appropriate for pessimistic. Described below are some examplepolicies.

Since each transactional memory (TM) system needs both versioningdetection and conflict detection, these options give rise to fourdistinct TM designs: Eager-Pessimistic (EP), Eager-Optimistic (EO),Lazy-Pessimistic (LP), and Lazy-Optimistic (LO). Table 2 brieflydescribes all four distinct TM designs.

FIG. 1 depicts an example of a multicore TM environment. FIG. 1 showsmany TM-enabled CPUs (CPU1 114 a, CPU2 114 b, etc.) on one die 100,connected with an interconnect 122, under management of an interconnectcontrol 120 a, 120 b. Each CPU 114 a, 114 b (also known as a Processor)may have a split cache consisting of an Instruction Cache 116 a, 116 bfor caching instructions from memory to be executed and a Data Cache 118a, 118 b with TM support for caching data (operands) of memory locationsto be operated on by CPU 114 a, 114 b (in FIG. 1, each CPU 114 a, 114 band its associated caches are referenced as 112 a, 112 b). In animplementation, caches of multiple dies 100 are interconnected tosupport cache coherency between the caches of the multiple dies 100. Inan implementation, a single cache, rather than the split cache isemployed holding both instructions and data. In implementations, the CPUcaches are one level of caching in a hierarchical cache structure. Forexample each die 100 may employ a shared cache 124 to be shared amongstall the CPUs on the die 100. In another implementation, each die mayhave access to a shared cache 124, shared amongst all the processors ofall the dies 100.

FIG. 2 shows the details of an example transactional CPU environment 112c, having a CPU 114 c, including additions to support TM. Thetransactional CPU (processor) 114 c may include hardware for supportingRegister Checkpoints 126 and special TM Registers 128. The transactionalCPU cache may have the MESI bits 130, Tags 140 and Data 142 of aconventional cache but also, for example, R bits 132 showing a line hasbeen read by the CPU 114 c while executing a transaction and W bits 138showing a line has been written-to by the CPU 114 c while executing atransaction.

A key detail for programmers in any TM system is how non-transactionalaccesses interact with transactions. By design, transactional accessesare screened from each other using the mechanisms above. However, theinteraction between a regular, non-transactional load with a transactioncontaining a new value for that address must still be considered. Inaddition, the interaction between a non-transactional store with atransaction that has read that address must also be explored. These areissues of the database concept isolation.

A TM system is said to implement strong isolation, sometimes calledstrong atomicity, when every non-transactional load and store acts likean atomic transaction. Therefore, non-transactional loads cannot seeuncommitted data and non-transactional stores cause atomicity violationsin any transactions that have read that address. A system where this isnot the case is said to implement weak isolation, sometimes called weakatomicity.

Strong isolation is often more desirable than weak isolation due to therelative ease of conceptualization and implementation of strongisolation. Additionally, if a programmer has forgotten to surround someshared memory references with transactions, causing bugs, then withstrong isolation, the programmer will often detect that oversight usinga simple debug interface because the programmer will see anon-transactional region causing atomicity violations. Also, programswritten in one model may work differently on another model.

Further, strong isolation is often easier to support in hardware TM thanweak isolation. With strong isolation, since the coherence protocolalready manages load and store communication between processors,transactions can detect non-transactional loads and stores and actappropriately. To implement strong isolation in software TransactionalMemory (TM), non-transactional code must be modified to include read-and write-barriers; potentially crippling performance. Although greateffort has been expended to remove many un-needed barriers, suchtechniques are often complex and performance is typically far lower thanthat of HTMs.

TABLE 2 Transactional Memory Design Space VERSIONING Lazy Eager CONFLICTOptimistic Storing updates in a write Not practical: waiting to updateDETECTION buffer; detecting conflicts at memory until commit time butcommit time. detecting conflicts at access time guarantees wasted workand provides no advantage Pessimistic Storing updates in a writeUpdating memory, keeping old buffer; detecting conflicts at values inundo log; detecting access time. conflicts at access time.

Table 2 illustrates the fundamental design space of transactional memory(versioning and conflict detection).

Eager-Pessimistic (EP)

This first TM design described below is known as Eager-Pessimistic. AnEP system stores its write-set “in place” (hence the name “eager”) and,to support rollback, stores the old values of overwritten lines in an“undo log”. Processors use the W 138 and R 132 cache bits to track readand write-sets and detect conflicts when receiving snooped loadrequests. Perhaps the most notable examples of EP systems in knownliterature are LogTM and UTM.

Beginning a transaction in an EP system is much like beginning atransaction in other systems: tm_begin( ) takes a register checkpoint,and initializes any status registers. An EP system also requiresinitializing the undo log, the details of which are dependent on the logformat, but often involve initializing a log base pointer to a region ofpre-allocated, thread-private memory, and clearing a log boundsregister.

Versioning: In EP, due to the way eager versioning is designed tofunction, the MESI 130 state transitions (cache line indicatorscorresponding to Modified, Exclusive, Shared, and Invalid code states)are left mostly unchanged. Outside of a transaction, the MESI 130 statetransitions are left completely unchanged. When reading a line inside atransaction, the standard coherence transitions apply (S (Shared)→S, I(Invalid)→S, or I→E (Exclusive)), issuing a load miss as needed, but theR 132 bit is also set. Likewise, writing a line applies the standardtransitions (S→M, E→I, I→M), issuing a miss as needed, but also sets theW 138 (Written) bit. The first time a line is written, the old versionof the entire line is loaded then written to the undo log to preserve itin case the current transaction aborts. The newly written data is thenstored “in-place,” over the old data.

Conflict Detection: Pessimistic conflict detection uses coherencemessages exchanged on misses, or upgrades, to look for conflicts betweentransactions. When a read miss occurs within a transaction, otherprocessors receive a load request; but they ignore the request if theydo not have the needed line. If the other processors have the neededline non-speculatively or have the line R 132 (Read), they downgradethat line to S, and in certain cases issue a cache-to-cache transfer ifthey have the line in MESI's 130 M or E state. However, if the cache hasthe line W 138, then a conflict is detected between the two transactionsand additional action(s) must be taken.

Similarly, when a transaction seeks to upgrade a line from shared tomodified (on a first write), the transaction issues an exclusive loadrequest, which is also used to detect conflicts. If a receiving cachehas the line non-speculatively, then the line is invalidated, and incertain cases a cache-to-cache transfer (M or E states) is issued. But,if the line is R 132 or W 138, a conflict is detected.

Validation: Because conflict detection is performed on every load, atransaction always has exclusive access to its own write-set. Therefore,validation does not require any additional work.

Commit: Since eager versioning stores the new version of data items inplace, the commit process simply clears the W 138 and R 132 bits anddiscards the undo log.

Abort: When a transaction rolls back, the original version of each cacheline in the undo log must be restored, a process called “unrolling” or“applying” the log. This is done during tm_discard( ) and must be atomicwith regard to other transactions. Specifically, the write-set muststill be used to detect conflicts: this transaction has the only correctversion of lines in its undo log, and requesting transactions must waitfor the correct version to be restored from that log. Such a log can beapplied using a hardware state machine or software abort handler.

Eager-Pessimistic has the characteristics of: Commit is simple and sinceit is in-place, very fast. Similarly, validation is a no-op. Pessimisticconflict detection detects conflicts early, thereby reducing the numberof “doomed” transactions. For example, if two transactions are involvedin a Write-After-Read dependency, then that dependency is detectedimmediately in pessimistic conflict detection. However, in optimisticconflict detection such conflicts are not detected until the writercommits.

Eager-Pessimistic also has the characteristics of: As described above,the first time a cache line is written, the old value must be written tothe log, incurring extra cache accesses. Aborts are expensive as theyrequire undoing the log. For each cache line in the log, a load must beissued, perhaps going as far as main memory before continuing to thenext line. Pessimistic conflict detection also prevents certainserializable schedules from existing.

Additionally, because conflicts are handled as they occur, there is apotential for livelock and careful contention management mechanisms mustbe employed to guarantee forward progress.

Lazy-Optimistic (LO)

Another popular TM design is Lazy-Optimistic (LO), which stores itswrite-set in a “write buffer” or “redo log” and detects conflicts atcommit time (still using the R 132 and W 138 bits).

Versioning: Just as in the EP system, the MESI protocol of the LO designis enforced outside of the transactions. Once inside a transaction,reading a line incurs the standard MESI transitions but also sets the R132 bit. Likewise, writing a line sets the W 138 bit of the line, buthandling the MESI transitions of the LO design is different from that ofthe EP design. First, with lazy versioning, the new versions of writtendata are stored in the cache hierarchy until commit while othertransactions have access to old versions available in memory or othercaches. To make available the old versions, dirty lines (M lines) mustbe evicted when first written by a transaction. Second, no upgrademisses are needed because of the optimistic conflict detection feature:if a transaction has a line in the S state, it can simply write to itand upgrade that line to an M state without communicating the changeswith other transactions because conflict detection is done at committime.

Conflict Detection and Validation: To validate a transaction and detectconflicts, LO communicates the addresses of speculatively modified linesto other transactions only when it is preparing to commit. Onvalidation, the processor sends one, potentially large, network packetcontaining all the addresses in the write-set. Data is not sent, butleft in the cache of the committer and marked dirty (M). To build thispacket without searching the cache for lines marked W, a simple bitvector is used, called a “store buffer,” with one bit per cache line totrack these speculatively modified lines. Other transactions use thisaddress packet to detect conflicts: if an address is found in the cacheand the R 132 and/or W 138 bits are set, then a conflict is initiated.If the line is found but neither R 132 nor W 138 is set, then the lineis simply invalidated, which is similar to processing an exclusive load.

To support transaction atomicity, these address packets must be handledatomically, i.e., no two address packets may exist at once with the sameaddresses. In an LO system, this can be achieved by simply acquiring aglobal commit token before sending the address packet. However, atwo-phase commit scheme could be employed by first sending out theaddress packet, collecting responses, enforcing an ordering protocol(perhaps oldest transaction first), and committing once all responsesare satisfactory.

Commit: Once validation has occurred, commit needs no special treatment:simply clear W 138 and R 132 bits and the store buffer. Thetransaction's writes are already marked dirty in the cache and othercaches' copies of these lines have been invalidated via the addresspacket. Other processors can then access the committed data through theregular coherence protocol.

Abort: Rollback is equally easy: because the write-set is containedwithin the local caches, these lines can be invalidated, then clear W138 and R 132 bits and the store buffer. The store buffer allows W linesto be found to invalidate without the need to search the cache.

Lazy-Optimistic has the characteristics of: Aborts are very fast,requiring no additional loads or stores and making only local changes.More serializable schedules can exist than found in EP, which allows anLO system to more aggressively speculate that transactions areindependent, which can yield higher performance. Finally, the latedetection of conflicts can increase the likelihood of forward progress.

Lazy-Optimistic also has the characteristics of: Validation takes globalcommunication time proportional to size of write set. Doomedtransactions can waste work since conflicts are detected only at committime.

Lazy-Pessimistic (LP)

Lazy-Pessimistic (LP) represents a third TM design option, sittingsomewhere between EP and LO: storing newly written lines in a writebuffer but detecting conflicts on a per access basis.

Versioning: Versioning is similar but not identical to that of LO:reading a line sets its R bit 132, writing a line sets its W bit 138,and a store buffer is used to track W lines in the cache. Also, dirty(M) lines must be evicted when first written by a transaction, just asin LO. However, since conflict detection is pessimistic, load exclusivesmust be performed when upgrading a transactional line from I, S→M, whichis unlike LO.

Conflict Detection: LP's conflict detection operates the same as EP's:using coherence messages to look for conflicts between transactions.

Validation: Like in EP, pessimistic conflict detection ensures that atany point, a running transaction has no conflicts with any other runningtransaction, so validation is a no-op.

Commit: Commit needs no special treatment: simply clear W 138 and R 132bits and the store buffer, like in LO.

Abort: Rollback is also like that of LO: simply invalidate the write-setusing the store buffer and clear the W and R bits and the store buffer.

Eager-Optimistic (EO)

The LP has the characteristics of: Like LO, aborts are very fast. LikeEP, the use of pessimistic conflict detection reduces the number of“doomed” transactions. Like EP, some serializable schedules are notallowed and conflict detection must be performed on each cache miss.

The final combination of versioning and conflict detection isEager-Optimistic (EO). EO may be a less than optimal choice for HTMsystems: since new transactional versions are written in-place, othertransactions have no choice but to notice conflicts as they occur (i.e.,as cache misses occur). But since EO waits until commit time to detectconflicts, those transactions become “zombies,” continuing to execute,wasting resources, yet are “doomed” to abort.

EO has proven to be useful in STMs and is implemented by Bartok-STM andMcRT. A lazy versioning STM needs to check its write buffer on each readto ensure that it is reading the most recent value. Since the writebuffer is not a hardware structure, this is expensive, hence thepreference for write-in-place eager versioning. Additionally, sincechecking for conflicts is also expensive in an STM, optimistic conflictdetection offers the advantage of performing this operation in bulk.

Contention Management

How a transaction rolls back once the system has decided to abort thattransaction has been described above, but, since a conflict involves twotransactions, the topics of which transaction should abort, how thatabort should be initiated, and when should the aborted transaction beretried need to be explored. These are topics that are addressed byContention Management (CM), a key component of transactional memory.Described below are policies regarding how the systems initiate abortsand the various established methods of managing which transactionsshould abort in a conflict.

Contention Management Policies

A Contention Management (CM) Policy is a mechanism that determines whichtransaction involved in a conflict should abort and when the abortedtransaction should be retried. For example, it is often the case thatretrying an aborted transaction immediately does not lead to the bestperformance. Conversely, employing a back-off mechanism, which delaysthe retrying of an aborted transaction, can yield better performance.STMs first grappled with finding the best contention management policiesand many of the policies outlined below were originally developed forSTMs.

CM Policies draw on a number of measures to make decisions, includingages of the transactions, size of read- and write-sets, the number ofprevious aborts, etc. The combinations of measures to make suchdecisions are endless, but certain combinations are described below,roughly in order of increasing complexity.

To establish some nomenclature, first note that in a conflict there aretwo sides: the attacker and the defender. The attacker is thetransaction requesting access to a shared memory location. Inpessimistic conflict detection, the attacker is the transaction issuingthe load or load exclusive. In optimistic, the attacker is thetransaction attempting to validate. The defender in both cases is thetransaction receiving the attacker's request.

An Aggressive CM Policy immediately and always retries either theattacker or the defender. In LO, Aggressive means that the attackeralways wins, and so Aggressive is sometimes called committer wins. Sucha policy was used for the earliest LO systems. In the case of EP,Aggressive can be either defender wins or attacker wins.

Restarting a conflicting transaction that will immediately experienceanother conflict is bound to waste work—namely interconnect bandwidthrefilling cache misses. A Polite CM Policy employs exponential backoff(but linear could also be used) before restarting conflicts. To preventstarvation, a situation where a process does not have resourcesallocated to it by the scheduler, the exponential backoff greatlyincreases the odds of transaction success after some n retries.

Another approach to conflict resolution is to randomly abort theattacker or defender (a policy called Randomized). Such a policy may becombined with a randomized backoff scheme to avoid unneeded contention.

However, making random choices, when selecting a transaction to abort,can result in aborting transactions that have completed “a lot of work”,which can waste resources. To avoid such waste, the amount of workcompleted on the transaction can be taken into account when determiningwhich transaction to abort. One measure of work could be a transaction'sage. Other methods include Oldest, Bulk TM, Size Matters, Karma, andPolka. Oldest is a simple timestamp method that aborts the youngertransaction in a conflict. Bulk TM uses this scheme. Size Matters islike Oldest but instead of transaction age, the number of read/writtenwords is used as the priority, reverting to Oldest after a fixed numberof aborts. Karma is similar, using the size of the write-set aspriority. Rollback then proceeds after backing off a fixed amount oftime. Aborted transactions keep their priorities after being aborted(hence the name Karma). Polka works like Karma but instead of backingoff a predefined amount of time, it backs off exponentially more eachtime.

Since aborting wastes work, it is logical to argue that stalling anattacker until the defender has finished their transaction would lead tobetter performance. Unfortunately, such a simple scheme easily leads todeadlock.

Deadlock avoidance techniques can be used to solve this problem. Greedyuses two rules to avoid deadlock. The first rule is, if a firsttransaction, T1, has lower priority than a second transaction, T0, or ifT1 is waiting for another transaction, then T1 aborts when conflictingwith T0. The second rule is, if T1 has higher priority than T0 and isnot waiting, then T0 waits until T1 commits, aborts, or starts waiting(in which case the first rule is applied). Greedy provides someguarantees about time bounds for executing a set of transactions. One EPdesign (LogTM) uses a CM policy similar to Greedy to achieve stallingwith conservative deadlock avoidance.

Example MESI coherency rules provide for four possible states in which acache line of a multiprocessor cache system may reside, M, E, S, and I,defined as follows:

Modified (M): The cache line is present only in the current cache, andis dirty; it has been modified from the value in main memory. The cacheis required to write the data back to main memory at some time in thefuture, before permitting any other read of the (no longer valid) mainmemory state. The write-back changes the line to the Exclusive state.

Exclusive (E): The cache line is present only in the current cache, butis clean; it matches main memory. It may be changed to the Shared stateat any time, in response to a read request. Alternatively, it may bechanged to the Modified state when writing to it.

Shared (S): Indicates that this cache line may be stored in other cachesof the machine and is “clean”; it matches the main memory. The line maybe discarded (changed to the Invalid state) at any time.

Invalid (I): Indicates that this cache line is invalid (unused).

TM coherency status indicators (R 132, W 138) may be provided for eachcache line, in addition to, or encoded in the MESI coherency bits. An R132 indicator indicates the current transaction has read from the dataof the cache line, and a W 138 indicator indicates the currenttransaction has written to the data of the cache line.

In another aspect of TM design, a system is designed using transactionalstore buffers. U.S. Pat. No. 6,349,361 titled “Methods and Apparatus forReordering and Renaming Memory References in a Multiprocessor ComputerSystem,” filed Mar. 31, 2000 and incorporated by reference herein in itsentirety, teaches a method for reordering and renaming memory referencesin a multiprocessor computer system having at least a first and a secondprocessor. The first processor has a first private cache and a firstbuffer, and the second processor has a second private cache and a secondbuffer. The method includes the steps of, for each of a plurality ofgated store requests received by the first processor to store a datum,exclusively acquiring a cache line that contains the datum by the firstprivate cache, and storing the datum in the first buffer. Upon the firstbuffer receiving a load request from the first processor to load aparticular datum, the particular datum is provided to the firstprocessor from among the data stored in the first buffer based on anin-order sequence of load and store operations. Upon the first cachereceiving a load request from the second cache for a given datum, anerror condition is indicated and a current state of at least one of theprocessors is reset to an earlier state when the load request for thegiven datum corresponds to the data stored in the first buffer.

The main implementation components of one such transactional memoryfacility are a transaction-backup register file for holdingpre-transaction GR (general register) content, a cache directory totrack the cache lines accessed during the transaction, a store cache tobuffer stores until the transaction ends, and firmware routines toperform various complex functions. In this section a detailedimplementation is described.

IBM zEnterprise EC12 Enterprise Server Embodiment

The IBM zEnterprise EC12 enterprise server introduces transactionalexecution (TX) in transactional memory, and is described in part in apaper, “Transactional Memory Architecture and Implementation for IBMSystem z” of Proceedings Pages 25-36 presented at MICRO-45, 1-5 Dec.2012, Vancouver, British Columbia, Canada, available from IEEE ComputerSociety Conference Publishing Services (CPS), which is incorporated byreference herein in its entirety.

Table 3 shows an example transaction. Transactions started with TBEGINare not assured to ever successfully complete with TEND, since they canexperience an aborting condition at every attempted execution, e.g., dueto repeating conflicts with other CPUs. This requires that the programsupport a fallback path to perform the same operationnon-transactionally, e.g., by using traditional locking schemes. Thisputs a significant burden on the programming and software verificationteams, especially where the fallback path is not automatically generatedby a reliable compiler.

TABLE 3 Example Transaction Code LHI R0,0 *initialize retry count = 0loop TBEGIN *begin transaction JNZ abort *go to abort code if CC1 = 0 LTR1, lock *load and test the fallback lock JNZ lckbzy *branch if lockbusy . . . perform operation . . . TEND *end transaction . . . . . . . .. . . . lckbzy TABORT *abort if lock busy; this *resumes after TBEGINabort JO fallback *no retry if CC = 3 AHI R0, 1 *increment retry countCIJNL R0,6, fallback *give up after 6 attempts PPA R0, TX *random delaybased on retry count . . . potentially wait for lock to become free . .. J loop *jump back to retry fallback OBTAIN lock *using Compare&Swap .. . perform operation . . . RELEASE lock . . . . . . . . . . . .

The requirement of providing a fallback path for aborted TransactionExecution (TX) transactions can be onerous. Many transactions operatingon shared data structures are expected to be short, touch only a fewdistinct memory locations, and use simple instructions only. For thosetransactions, the IBM zEnterprise EC12 introduces the concept ofconstrained transactions; under normal conditions, the CPU 114 c(FIG. 1) assures that constrained transactions eventually endsuccessfully, albeit without giving a strict limit on the number ofnecessary retries. A constrained transaction starts with a TBEGINCinstruction and ends with a regular TEND. Implementing a task as aconstrained or non-constrained transaction typically results in verycomparable performance, but constrained transactions simplify softwaredevelopment by removing the need for a fallback path. IBM'sTransactional Execution architecture is further described inz/Architecture, Principles of Operation, Tenth Edition, SA22-7832-09published September 2012 from IBM, incorporated by reference herein inits entirety.

A constrained transaction starts with the TBEGINC instruction. Atransaction initiated with TBEGINC must follow a list of programmingconstraints; otherwise the program takes a non-filterableconstraint-violation interruption. Exemplary constraints may include,but not be limited to: the transaction can execute a maximum of 32instructions, all instruction text must be within 256 consecutive bytesof memory; the transaction contains only forward-pointing relativebranches (i.e., no loops or subroutine calls); the transaction canaccess a maximum of 4 aligned octowords (an octoword is 32 bytes) ofmemory; and restriction of the instruction-set to exclude complexinstructions like decimal or floating-point operations. The constraintsare chosen such that many common operations like doubly linkedlist-insert/delete operations can be performed, including the verypowerful concept of atomic compare-and-swap targeting up to 4 alignedoctowords. At the same time, the constraints were chosen conservativelysuch that future CPU implementations can assure transaction successwithout needing to adjust the constraints, since that would otherwiselead to software incompatibility.

TBEGINC mostly behaves like XBEGIN in TSX or TBEGIN on IBM's zEC12servers, except that the floating-point register (FPR) control and theprogram interruption filtering fields do not exist and the controls areconsidered to be zero. On a transaction abort, the instruction addressis set back directly to the TBEGINC instead of to the instruction after,reflecting the immediate retry and absence of an abort path forconstrained transactions.

Nested transactions are not allowed within constrained transactions, butif a TBEGINC occurs within a non-constrained transaction it is treatedas opening a new non-constrained nesting level just like TBEGIN would.This can occur, e.g., if a non-constrained transaction calls asubroutine that uses a constrained transaction internally.

Since interruption filtering is implicitly off, all exceptions during aconstrained transaction lead to an interruption into the operatingsystem (OS). Eventual successful finishing of the transaction relies onthe capability of the OS to page-in the at most 4 pages touched by anyconstrained transaction. The OS must also ensure time-slices long enoughto allow the transaction to complete.

TABLE 4 Transaction Code Example TBEGINC *begin constrained transaction. . . perform operation . . . TEND *end transaction

Table 4 shows the constrained-transactional implementation of the codein Table 3, assuming that the constrained transactions do not interactwith other locking-based code. No lock testing is shown therefore, butcould be added if constrained transactions and lock-based code weremixed.

When failure occurs repeatedly, software emulation is performed usingmillicode as part of system firmware. Advantageously, constrainedtransactions have desirable properties because of the burden removedfrom programmers.

With reference to FIG. 2, the IBM zEnterprise EC12 processor introducedthe transactional execution facility. The processor can decode 3instructions per clock cycle; simple instructions are dispatched assingle micro-ops, and more complex instructions are cracked intomultiple micro-ops. The micro-ops (uops 232 b) are written into aunified issue queue 216, from where they can be issued out-of-order. Upto two fixed-point, one floating-point, two load/store, and two branchinstructions can execute every cycle. A Global Completion Table (GCT)232 holds every micro-op 232 b and a transaction nesting depth (TND) 232a. The GCT 232 is written in-order at decode time, tracks the executionstatus of each micro-op 232 b, and completes instructions when allmicro-ops 232 b of the oldest instruction group have successfullyexecuted.

The level 1 (L1) data cache 240 is a 96 KB (kilo-byte) 6-way associativecache with 256 byte cache-lines and 4 cycle use latency, coupled to aprivate 1 MB (mega-byte) 8-way associative 2nd-level (L2) data cache 268with 7 cycles use-latency penalty for L1 240 misses. The L1 240 cache isthe cache closest to a processor and Ln cache is a cache at the nthlevel of caching. Both L1 240 and L2 268 caches are store-through. Sixcores on each central processor (CP) chip share a 48 MB 3rd-levelstore-in cache, and six CP chips are connected to an off-chip 384 MB4th-level cache, packaged together on a glass ceramic multi-chip module(MCM). Up to 4 multi-chip modules (MCMs) can be connected to a coherentsymmetric multi-processor (SMP) system with up to 144 cores (not allcores are available to run customer workload).

Coherency is managed with a variant of the MESI protocol. Cache-linescan be owned read-only (shared) or exclusive; the L1 240 and L2 268 arestore-through and thus do not contain dirty lines. The L3 272 and L4caches (not shown) are store-in and track dirty states. Each cache isinclusive of all its connected lower level caches.

Coherency requests are called “cross interrogates” (XI) and are senthierarchically from higher level to lower-level caches, and between theL4s. When one core misses the L1 240 and L2 268 and requests the cacheline from its local L3 272, the L3 272 checks whether it owns the line,and if necessary sends an XI to the currently owning L2 268/L1 240 underthat L3 272 to ensure coherency, before it returns the cache line to therequestor. If the request also misses the L3 272, the L3 272 sends arequest to the L4 (not shown), which enforces coherency by sending XIsto all necessary L3s under that L4, and to the neighboring L4s. Then theL4 responds to the requesting L3 which forwards the response to the L2268/L1 240.

Note that due to the inclusivity rule of the cache hierarchy, sometimescache lines are XI'ed from lower-level caches due to evictions onhigher-level caches caused by associativity overflows from requests toother cache lines. These XIs can be called “LRU XIs”, where LRU standsfor least recently used.

Making reference to yet another type of XI requests, Demote-XIstransition cache-ownership from exclusive into read-only state, andExclusive-XIs transition cache ownership from exclusive into invalidstate. Demote-XIs and Exclusive-XIs need a response back to the XIsender. The target cache can “accept” the XI, or send a “reject”response if it first needs to evict dirty data before accepting the XI.The L1 240/L2 268 caches are store through, but may reject demote-XIsand exclusive XIs if they have stores in their store queues that need tobe sent to L3 before downgrading the exclusive state. A rejected XI willbe repeated by the sender. Read-only-XIs are sent to caches that own theline read-only; no response is needed for such XIs since they cannot berejected. The details of the SMP protocol are similar to those describedfor the IBM z10 by P. Mak, C. Walters, and G. Strait, in “IBM System z10processor cache subsystem microarchitecture”, IBM Journal of Researchand Development, Vol 53:1, 2009, which is incorporated by referenceherein in its entirety.

Transactional Instruction Execution

FIG. 2 depicts example components of an example transactional executionenvironment, including a CPU and caches/components with which itinteracts (such as those depicted in FIG. 1). The instruction decodeunit 208 (IDU) keeps track of the current transaction nesting depth 212(TND). When the IDU 208 receives a TBEGIN instruction, the nesting depth212 is incremented, and conversely decremented on TEND instructions. Thenesting depth 212 is written into the GCT 232 for every dispatchedinstruction. When a TBEGIN or TEND is decoded on a speculative path thatlater gets flushed, the IDU's 208 nesting depth 212 is refreshed fromthe youngest GCT 232 entry that is not flushed. The transactional stateis also written into the issue queue 216 for consumption by theexecution units, mostly by the Load/Store Unit (LSU) 280, which also hasan effective address calculator 236 included in the LSU 280. The TBEGINinstruction may specify a transaction diagnostic block (TDB) forrecording status information, should the transaction abort beforereaching a TEND instruction.

Similar to the nesting depth, the IDU 208/GCT 232 collaboratively trackthe access register/floating-point register (AR/FPR) modification masksthrough the transaction nest; the IDU 208 can place an abort requestinto the GCT 232 when an AR/FPR-modifying instruction is decoded and themodification mask blocks that. When the instruction becomesnext-to-complete, completion is blocked and the transaction aborts.Other restricted instructions are handled similarly, including TBEGIN ifdecoded while in a constrained transaction, or exceeding the maximumnesting depth.

An outermost TBEGIN is cracked into multiple micro-ops depending on theGR-Save-Mask; each micro-op 232 b (including, for example uop 0, uop 1,and uop 2) will be executed by one of the two fixed point units (FXUs)220 to save a pair of GRs 228 into a special transaction-backup registerfile 224, that is used to later restore the GR 228 content in case of atransaction abort. Also the TBEGIN spawns micro-ops 232 b to perform anaccessibility test for the TDB if one is specified; the address is savedin a special purpose register for later usage in the abort case. At thedecoding of an outermost TBEGIN, the instruction address and theinstruction text of the TBEGIN are also saved in special purposeregisters for a potential abort processing later on.

TEND and NTSTG are single micro-op 232 b instructions; NTSTG(non-transactional store) is handled like a normal store except that itis marked as non-transactional in the issue queue 216 so that the LSU280 can treat it appropriately. TEND is a no-op at execution time, theending of the transaction is performed when TEND completes.

As mentioned, instructions that are within a transaction are marked assuch in the issue queue 216, but otherwise execute mostly unchanged; theLSU 280 performs isolation tracking as described in the next section.

Since decoding is in-order, and since the IDU 208 keeps track of thecurrent transactional state and writes it into the issue queue 216 alongwith every instruction from the transaction, execution of TBEGIN, TEND,and instructions before, within, and after the transaction can beperformed out-of order. It is even possible (though unlikely) that TENDis executed first, then the entire transaction, and lastly the TBEGINexecutes. Program order is restored through the GCT 232 at completiontime. The length of transactions is not limited by the size of the GCT232, since general purpose registers (GRs) 228 can be restored from thebackup register file 224.

During execution, the program event recording (PER) events are filteredbased on the Event Suppression Control, and a PER TEND event is detectedif enabled. Similarly, while in transactional mode, a pseudo-randomgenerator may be causing the random aborts as enabled by the TransactionDiagnostics Control.

Tracking for Transactional Isolation

The Load/Store Unit 280 tracks cache lines that were accessed duringtransactional execution, and triggers an abort if an XI from another CPU(or an LRU-XI) conflicts with the footprint. If the conflicting XI is anexclusive or demote XI, the LSU 280 rejects the XI back to the L3 272 inthe hope of finishing the transaction before the L3 272 repeats the XI.This “stiff-arming” is very efficient in highly contended transactions.In order to prevent hangs when two CPUs stiff-arm each other, aXI-reject counter is implemented, which triggers a transaction abortwhen a threshold is met.

The L1 cache directory 240 is traditionally implemented with staticrandom access memories (SRAMs). For the transactional memoryimplementation, the valid bits 244 (64 rows×6 ways) of the directoryhave been moved into normal logic latches, and are supplemented with twomore bits per cache line: the TX-read 248 and TX-dirty 252 bits.

The TX-read 248 bits are reset when a new outermost TBEGIN is decoded(which is interlocked against a prior still pending transaction). TheTX-read 248 bit is set at execution time by every load instruction thatis marked “transactional” in the issue queue. Note that this can lead toover-marking if speculative loads are executed, for example on amispredicted branch path. The alternative of setting the TX-read 248 bitat load completion time was too expensive for silicon area, sincemultiple loads can complete at the same time, requiring many read-portson the load-queue.

Stores execute the same way as in non-transactional mode, but atransaction mark is placed in the store queue (STQ) 260 entry of thestore instruction. At write-back time, when the data from the STQ 260 iswritten into the L1 240, the TX-dirty bit 252 in the L1-directory 256 isset for the written cache line. Store write-back into the L1 240 occursonly after the store instruction has completed, and at most one store iswritten back per cycle. Before completion and write-back, loads canaccess the data from the STQ 260 by means of store-forwarding; afterwrite-back, the CPU 114 c (FIG. 1) can access the speculatively updateddata in the L1 240. If the transaction ends successfully, the TX-dirtybits 252 of all cache-lines are cleared, and also the TX-marks of notyet written stores are cleared in the STQ 260, effectively turning thepending stores into normal stores.

On a transaction abort, all pending transactional stores are invalidatedfrom the STQ 260, even those already completed. All cache lines thatwere modified by the transaction in the L1 240, that is, have theTX-dirty bit 252 on, have their valid bits turned off, effectivelyremoving them from the L1 240 cache instantaneously.

The architecture requires that before completing a new instruction, theisolation of the transaction read- and write-set is maintained. Thisisolation is ensured by stalling instruction completion at appropriatetimes when XIs are pending; speculative out-of order execution isallowed, optimistically assuming that the pending XIs are to differentaddresses and not actually cause a transaction conflict. This designfits very naturally with the XI-vs-completion interlocks that areimplemented on prior systems to ensure the strong memory ordering thatthe architecture requires.

When the L1 240 receives an XI, L1 240 accesses the directory to checkvalidity of the XI′ed address in the L1 240, and if the TX-read bit 248is active on the XI′ed line and the XI is not rejected, the LSU 280triggers an abort. When a cache line with active TX-read bit 248 isLRU′ed from the L1 240, a special LRU-extension vector remembers foreach of the 64 rows of the L1 240 that a TX-read line existed on thatrow. Since no precise address tracking exists for the LRU extensions,any non-rejected XI that hits a valid extension row the LSU 280 triggersan abort. Providing the LRU-extension effectively increases the readfootprint capability from the L1-size to the L2-size and associativity,provided no conflicts with other CPUs 114 (FIG. 1) against thenon-precise LRU-extension tracking causes aborts.

The store footprint is limited by the store cache size (the store cacheis discussed in more detail below) and thus implicitly by the L2 268size and associativity. No LRU-extension action needs to be performedwhen a TX-dirty 252 cache line is LRU′ed from the L1 240.

Store Cache

In prior systems, since the L1 240 and L2 268 are store-through caches,every store instruction causes an L3 272 store access; with now 6 coresper L3 272 and further improved performance of each core, the store ratefor the L3 272 (and to a lesser extent for the L2 268) becomesproblematic for certain workloads. In order to avoid store queuingdelays, a gathering store cache 264 had to be added, that combinesstores to neighboring addresses before sending them to the L3 272.

For transactional memory performance, it is acceptable to invalidateevery TX-dirty 252 cache line from the L1 240 on transaction aborts,because the L2 268 cache is very close (7 cycles L1 240 miss penalty) tobring back the clean lines. However, it would be unacceptable forperformance (and silicon area for tracking) to have transactional storeswrite the L2 268 before the transaction ends and then invalidate alldirty L2 268 cache lines on abort (or even worse on the shared L3 272).

The two problems of store bandwidth and transactional memory storehandling can both be addressed with the gathering store cache 264. Thecache 264 is a circular queue of 64 entries, each entry holding 128bytes of data with byte-precise valid bits. In non-transactionaloperation, when a store is received from the LSU 280, the store cache264 checks whether an entry exists for the same address, and if sogathers the new store into the existing entry. If no entry exists, a newentry is written into the queue, and if the number of free entries fallsunder a threshold, the oldest entries are written back to the L2 268 andL3 272 caches.

When a new outermost transaction begins, all existing entries in thestore cache are marked closed so that no new stores can be gathered intothem, and eviction of those entries to L2 268 and L3 272 is started.From that point on, the transactional stores coming out of the LSU 280STQ 260 allocate new entries, or gather into existing transactionalentries. The write-back of those stores into L2 268 and L3 272 isblocked, until the transaction ends successfully; at that pointsubsequent (post-transaction) stores can continue to gather intoexisting entries, until the next transaction closes those entries again.

The store cache 264 is queried on every exclusive or demote XI, andcauses an XI reject if the XI compares to any active entry. If the coreis not completing further instructions while continuously rejecting XIs,the transaction is aborted at a certain threshold to avoid hangs.

The LSU 280 requests a transaction abort when the store cache 264overflows. The LSU 280 detects this condition when it tries to send anew store that cannot merge into an existing entry, and the entire storecache 264 is filled with stores from the current transaction. The storecache 264 is managed as a subset of the L2 268: while transactionallydirty lines can be evicted from the L1 240, they have to stay residentin the L2 268 throughout the transaction. The maximum store footprint isthus limited to the store cache size of 64×128 bytes, and it is alsolimited by the associativity of the L2 268. Since the L2 268 is 8-wayassociative and has 512 rows, it is typically large enough to not causetransaction aborts.

If a transaction aborts, the store cache 264 is notified and all entriesholding transactional data are invalidated. The store cache 264 also hasa mark per doubleword (8 bytes) whether the entry was written by a NTSTGinstruction—those doublewords stay valid across transaction aborts.

Millicode-Implemented Functions

Traditionally, IBM mainframe server processors contain a layer offirmware called millicode which performs complex functions like certainCISC instruction executions, interruption handling, systemsynchronization, and RAS. Millicode includes machine dependentinstructions as well as instructions of the instruction set architecture(ISA) that are fetched and executed from memory similarly toinstructions of application programs and the operating system (OS).Firmware resides in a restricted area of main memory that customerprograms cannot access. When hardware detects a situation that needs toinvoke millicode, the instruction fetching unit 204 switches into“millicode mode” and starts fetching at the appropriate location in themillicode memory area. Millicode may be fetched and executed in the sameway as instructions of the instruction set architecture (ISA), and mayinclude ISA instructions.

For transactional memory, millicode is involved in various complexsituations. Every transaction abort invokes a dedicated millicodesub-routine to perform the necessary abort steps. The transaction-abortmillicode starts by reading special-purpose registers (SPRs) holding thehardware internal abort reason, potential exception reasons, and theaborted instruction address, which millicode then uses to store a TDB ifone is specified. The TBEGIN instruction text is loaded from an SPR toobtain the GR-save-mask, which is needed for millicode to know which GRs238 to restore.

The CPU 114 c (FIG. 1) supports a special millicode-only instruction toread out the backup-GRs 224 and copy them into the main GRs 228. TheTBEGIN instruction address is also loaded from an SPR to set the newinstruction address in the PSW to continue execution after the TBEGINonce the millicode abort sub-routine finishes. That PSW may later besaved as program-old PSW in case the abort is caused by a non-filteredprogram interruption.

The TABORT instruction may be millicode implemented; when the IDU 208decodes TABORT, it instructs the instruction fetch unit to branch intoTABORT's millicode, from which millicode branches into the common abortsub-routine.

The Extract Transaction Nesting Depth (ETND) instruction may also bemillicoded, since it is not performance critical; millicode loads thecurrent nesting depth out of a special hardware register and places itinto a GR 228. The PPA instruction is millicoded; it performs theoptimal delay based on the current abort count provided by software asan operand to PPA, and also based on other hardware internal state.

For constrained transactions, millicode may keep track of the number ofaborts. The counter is reset to 0 on successful TEND completion, or ifan interruption into the OS occurs (since it is not known if or when theOS will return to the program). Depending on the current abort count,millicode can invoke certain mechanisms to improve the chance of successfor the subsequent transaction retry. The mechanisms involve, forexample, successively increasing random delays between retries, andreducing the amount of speculative execution to avoid encounteringaborts caused by speculative accesses to data that the transaction isnot actually using. As a last resort, millicode can broadcast to otherCPUs 114 (FIG. 1) to stop all conflicting work, retry the localtransaction, before releasing the other CPUs 114 to continue normalprocessing. Multiple CPUs 114 must be coordinated to not causedeadlocks, so some serialization between millicode instances ondifferent CPUs 114 is required.

Managing Transactional Aborts

As previously described, events occurring during execution of atransaction may disrupt the transaction and cause the transaction toabort. A disruptive event may be any event that can result in atransactional abort. For example, a disruptive event may be a systemevent, such as a pending timer interrupt or other pending asynchronousinterrupt. For another example, a disruptive event may be a conflictwith a second processor, such as when the second processor's activities,such as cache or other memory accesses, interfere with the transactioncode region's ability to execute transactionally.

When a processor executing a non-constrained transaction identifies sucha disruptive event, the processor has at least two possible actions: (1)the processor may immediately abort the transaction in favor of thedisruptive event, and (2) the processor may defer processing of thedisruptive event for some period of time and attempt to complete thetransaction, thus avoiding the performance loss associated with atransactional abort. Time and resources spent processing a transactionthat is eventually aborted are wasted, and if a transaction is abortedtoo many times, software may need to resort to a course-grained lock ortake other drastic measures in order to complete the functions attemptedby the transaction. Such drastic measures may result in even furtherperformance losses.

For example, if a processor executing a transaction identifies a pendingtimer interrupt, the processor may defer processing that timer interruptin favor of attempting to complete the transaction. For another example,if a processor executing a transaction identifies an incoming cachecoherency request involving the transaction's cache footprint, theprocessor may temporarily refuse to honor the coherency request in favorof attempting to complete the transaction. As previously described, acoherency request may be broadcast from a processor to gain exclusiveaccess to a cache line, such as when the processor needs to write to thecache line, or may result if a snoop of the cache line on a commonaddress request bus reveals an inconsistency with more recent data. Insome implementations, a coherency request may be referred to as across-interrogate (XI).

Ironically, a processor that defers processing of a disruptive event infavor of attempting to complete a transaction may actually becontributing to a loss in overall system performance. In the examplesabove, deferring the timer interrupt may degrade system responsiveness,and refusing to honor the coherency request may slow down overall systemprocessing, or in extreme cases may cause a system hang if deferredindefinitely. Furthermore, a processor's attempt to complete atransaction may still eventually fail since disruptive events may not bedeferred indefinitely. Time and resources spent on failed attempts tocomplete a transaction are doubly wasteful; in addition to the resourceswasted by continuing to process a doomed transaction, immediateprocessing of the deferred disruptive event would likely have improvedsystem performance.

In determining whether to defer processing of a disruptive event infavor of attempting to complete a transaction, it would be helpful forthe processor to know how close the transaction is to completion. Atransaction that is closer to completion may be a good candidate for anattempt to complete the transaction, while a transaction that is furtherfrom completion may be a good candidate for an immediate abort in favorof processing the disruptive event.

One way to provide the processor with an estimate of how close thetransaction is to completion is to introduce a new feature to thetransaction-begin instruction. Because different implementations may usedifferent terminology when referencing the instruction that indicatesthe start of a transaction (for example XBEGIN, XACQUIRE, or TBEGIN),this disclosure uses the term “transaction-begin instruction” or TBEGINwhen referencing the instruction that indicates the start of atransaction. The new feature of the transaction-begin instructionprovides a hint to the processor of approximately how long thetransaction will take to complete by specifying a value. This value isreferred to within this disclosure as a“transaction-count-to-completion” (CTC) value. The specified CTC valueis an approximation of how close the transaction is to completion at thetime the transaction-begin instruction is executed; in other words, thespecified CTC value is an approximation of how long the transaction willtake to complete. As the transaction progresses and more instructionsare executed, the processor may periodically adjust the internal CTCvalue, for example, the processor may adjust the CTC value after eachinstruction is executed. The adjusted CTC value accounts for theprogress of the transaction since execution of the transaction-begininstruction. If the processor identifies a disruptive event whileprocessing the transaction, the processor may use the adjusted CTC valueto determine whether to defer processing of the disruptive event orwhether to immediately abort the transaction in favor of the disruptiveevent.

In some embodiments, the initial CTC value specified by thetransaction-begin instruction may indicate a predicted number ofinstructions to execute in the transaction. In some embodiments, theinitial CTC value specified by the transaction-begin instruction mayindicate a predicted amount of time remaining to execute thetransaction. The transaction-begin instruction may specify the initialCTC value using zero or more instruction operands; such operands may beimmediate values, register locations, memory locations, or a combinationof these and other types of operands. A transaction-begin instructionwith zero operands for specifying the initial CTC value may specify theinitial CTC value in another way, for example by an implied associationwith a particular register and/or memory location upon which the initialCTC value may be based.

Especially for moderate to complex transactions, for exampletransactions that include branches, it may be difficult for atransaction-begin instruction to specify a precisely accurate initialCTC value. Consequently, the initial CTC value may often be anapproximation of the number of instructions required to complete thetransaction, or an approximation of the amount of time required tocomplete the transaction. A programmer or compiler may attempt toprovide an initial CTC value that covers most contingencies through thebranch paths of the transaction, may attempt to provide a “worst-case”CTC value that covers the longest path through the transaction, or mayprovide a CTC value based on some other criteria. For transactions wherea CTC value cannot be determined at compile time, such as whentransaction code follows a linked list pointer chain, a programmer orcompiler may provide an initial CTC value that is based on an averagevalue determined via experimental execution of the transaction code. Forexample, the initial CTC value provided may be an amount of time that issomewhat greater than the average time determined via experimentalexecution of the transaction code. In some embodiments, software maydynamically determine the time required for individual transactions tocomplete. Transaction-begin instructions for a previously timedtransaction may then specify that previously determined time as theinitial CTC value. For example, Java code may determine the timerequired for a transaction to complete before it JITs the code; theJITed code may then contain a CTC value while the interpreted code maynot.

An example system 300 of an embodiment for managing transactional abortswith a transaction-begin instruction is illustrated in FIG. 3. A memory315 contains instructions 325 for execution by a processor 305. Althoughinstructions 325 are represented as a contiguous block in a singlememory 315 in system 300, the format and location of instructions 325may vary among embodiments.

Within instructions 325 is a transaction 330. As previously described, atransaction is a programmer-specified code region in which most memoryoperations within the region appear to have occurred instantaneouslyupon successful completion. In the embodiment illustrated in system 300,a TBEGIN instruction 335 specifies the start of transaction 330 and aTEND instruction 340 specifies the end of transaction 330.

Besides specifying the start of transaction 330, the TBEGIN instruction335 provides a hint to the processor of approximately how long thattransaction will take to complete. In some embodiments, the TBEGINinstruction may also specify other values or provide other information.

When processor 305 executes TBEGIN instruction 335, it determines aninitial CTC value indicating approximately how long transaction 330 willtake to complete. In some embodiments, processor 305 may determine theinitial CTC value directly from TBEGIN instruction 335, for example, byinspecting an operand associated with TBEGIN instruction 335 thatcontains a compile-time constant. In some embodiments, processor 305 maydetermine the initial CTC value indirectly from TBEGIN instruction 335,for example, by inspecting a run-time value contained in a register,such as a register from CPU registers 310, or in a particular memorylocation, such as a memory location in data memory 320. Such register ormemory locations may be specified in operands associated with TBEGINinstruction 335. In some embodiments, processor 305 may determine theinitial CTC value based on a combination of compile-time constants andrun-time values associated with TBEGIN instruction 335.

As processor 305 executes instructions within transaction 330 afterprocessing TBEGIN instruction 335, the processor adjusts the CTC valuebased on the progress of the transaction. In some embodiments, processor305 may load an internal CTC counter with the specified CTC value fromTBEGIN instruction 335, and may adjust the CTC value by decrementing theinternal CTC counter as instructions are completed. In some embodiments,processor 305 may increment a counter to count instructions as they arecompleted, and may adjust the CTC value by first determining thedifference between the current value of the counter and the value at thetime TBEGIN instruction 335 was executed, and then subsequentlycomparing this determined difference to the specified initial CTC valuefrom TBEGIN instruction 335.

In some embodiments, processor 305 may use one or more timer constructsto adjust the CTC value. For example, processor 305 may start a timerwhen TBEGIN instruction 335 is executed, and may adjust the CTC value bysubtracting the timer value from the initial CTC value specified byTBEGIN instruction 335 as instructions are completed. For anotherexample, processor 305 may load an internal CPU register with thespecified initial CTC value in clock cycles from TBEGIN instruction 335,and may adjust the CTC value by decrementing the internal CPU registeras clock cycles elapse.

Various methods and processes for adjusting the CTC value arecontemplated, and any and all are appropriate so long as the adjustedCTC value accounts for the progress of the transaction since TBEGINinstruction 335 was executed. Processor 305 may then use a transaction'sadjusted CTC value in determining whether to defer processing of adisruptive event in favor of attempting to complete the transaction.

The format of a TBEGIN instruction may include an operation code 350,zero or more operands for use in specifying the initial CTC value, andzero or more operands 352 for uses unrelated to specifying the initialCTC value. Although system 300 shows an instruction format with threeoptional operands 354, 356, and 358 for use in specifying the initialCTC value, some embodiments may provide for exactly one operand, someembodiments may provide for exactly two operands, some embodiments mayprovide for one mandatory operand and one optional operand, and otherembodiments may provide for a different combination of mandatory andoptional operands. The operands may be defined as immediate valuefields, register designation fields, memory address fields, mask fields,length fields, displacement fields, or any other type of field,depending on the embodiment. Several non-limiting examples of TBEGINinstruction formats follow, and many more are contemplated.

The Instruction Set Architecture (ISA) may include a transaction-begininstruction that specifies an initial CTC value indicating theapproximate number of instructions remaining before the transactioncompletes. This transaction-begin instruction may have an operation code360 represented by the mnemonic TBEGIN, may require zero or moreoperands 361 unrelated to specifying the initial CTC value, and mayrequire one immediate-value operand 362 specifying the initial CTC valueas a number of instructions. For another example, the ISA may include atransaction-begin instruction that specifies an initial CTC valueindicating the approximate amount of time remaining before thetransaction completes. This transaction-begin instruction may have anoperation code 370 represented by the mnemonic TBEGIN, may require zeroor more operands 371 unrelated to specifying the initial CTC value, andmay require one immediate-value operand 372 specifying the initial CTCvalue as a number of clock cycles, picoseconds, nanoseconds,microseconds, or some other representation of time. For another example,a TBEGIN instruction 380 may require zero or more operands 381 unrelatedto specifying the initial CTC value, one immediate-value operand 382specifying a time value, and an optional operand 384 specifying a memoryaddress that contains a dynamically determined time value based onsoftware run-time history. Together, these two time values may specifythe initial CTC value, for example by adding them or subtracting onefrom the other. For another example, a TBEGIN instruction 390 may followthe RXY format in the z/Architecture, allowing for both a compile timeconstant (immediate value 392) and a dynamically determined value basedon software run-time history (register designation field 394) forspecifying the initial CTC value, and may allow for zero or moreoperands 391 unrelated to specifying the initial CTC value. For anotherexample, the ISA may include a generic transaction-begin instructionthat specifies an initial CTC value and further specifies whether theCTC value indicates the approximate number of instructions remainingbefore the transaction completes or the approximate amount of timeremaining before the transaction completes. This transaction-begininstruction may have an operation code 395 represented by the mnemonicTBEGIN, may require a mask operand 397 with an indicator flag specifyingwhether the initial CTC value represents a number of instructions or anamount of time, and may require one immediate-value operand 399specifying the CTC value, in addition to zero or more operands 396unrelated to specifying the initial CTC value. Example code containing aTBEGIN instruction with an instruction-based operand is shown below inTable 5.

TABLE 5 Example Transaction Code TBEGIN OTHER, 8(, 0) *8 remaininginstructions in TX BRC 7, ABORT_HDLR *Call software abort handler L 1, A*Fetch A and B atomically L 2, B CLR 1, 2 *Compare A and B BRE A_EQ_B*If A = B, we are done ST 1, B *Swap A and B atomically ST 2, A A_EQ_BTEND *End TX

In the Table 5 example, the transaction-begin instruction (TBEGIN)specifies a CTC value of 8, indicating that there are eight instructionsin the transaction. Note that if A equals B, there are actually only sixinstructions, but the programmer has selected a worst-case value of 8for the CTC value. Also note that an embodiment may include one, both,or neither of the TBEGIN and TEND instructions in the specified CTCvalue. As the processor executes each subsequent instruction, theprocessor adjusts the CTC value. After the BRC instruction, the CTCvalue is 7. After the first load (L) instruction, the CTC value is 6.After the second load (L) instruction, the CTC value is 5. If A does notequal B, the CTC value is 3 when the processor starts processing thestore (ST) instructions. If the processor then identifies a disruptiveevent, and if the adjusted CTC value of 3 satisfies deferral criteria,the processor may defer processing the disruptive event until A and Bare stored and the transaction is complete. The deferral criteria maybe, for example, a predefined threshold number determined by engineeringtesting for this processor and system environment. Deferral criteria maybe hard-coded, directly programmed, obtained through a user interface,determined at run-time, or supplied in some other manner.

An example method 400 of an embodiment for managing transactional abortswith a transaction-begin instruction is illustrated in FIG. 4. At start405, the processor may get the next instruction at 410. If the nextinstruction does not indicate the start of a transactional code regionat 415, then the processor may execute or otherwise process theinstruction at 420 and return to get the next instruction at 410.

If the next instruction indicates the start of a transactional coderegion at 415, then the processor may determine the initial CTC valuefrom the transaction-begin instruction at 445 and get the nextinstruction in the transaction at 425. If the next instruction in thetransaction indicates the end of the transactional code region at 430,then the instruction is executed at 432, the atomic commit issuccessful, and the transaction is complete. If the next instruction inthe transaction does not indicate the end of the transactional coderegion at 430, then the processor may execute or otherwise process theinstruction at 450. Although method 500 shows instruction execution at450, in some embodiments, instruction execution may be postponed untilafter deferment is evaluated, for example after 455, 460, 465, or 480.

After executing the instruction at 450, the processor may determinewhether a previously deferred disruptive event is pending at 455.Although in method 400 the processor checks for a pending deferreddisruptive event after each instruction in the transaction is executed,in some embodiments the processor may check more frequently or lessfrequently. The processor may also determine whether a previouslydeferred disruptive event is pending at 440 after the transaction iscomplete at 432. If a deferred disruptive event is pending at 440, afterthe transaction is complete, then the processor may process the deferreddisruptive event at 485 and return to get the next instruction at 410.If a deferred disruptive event is not pending at 440, the processor mayreturn to get the next instruction at 410.

If a deferred disruptive event is pending at 455, before the transactionis complete, then the processor may determine whether a deferral delayhas expired at 460. The deferral delay may represent the longestdeferral allowed in the environment, and may prevent a disruptive eventfrom being deferred for too long. The deferral delay may be defined as anumber of instructions, as a length of time, or in some other manner.The deferral delay may be, for example, a predefined threshold number ofinstructions or a predetermined number of clock cycles determined byengineering testing. The deferral delay may be hard-coded, directlyprogrammed, obtained through a user interface, determined at run-time,or supplied in some other manner. Determining whether the deferral delayhas expired may be accomplished through polling, interrupt handling, orin some other manner.

A deferral delay may be useful in dealing with malevolent software wherea programmer uses a transaction-begin instruction to specify anartificially low CTC value, for example in an attempt to defer abortsindefinitely. Expiration of a deferral delay may cause the processor toprocess a deferred disruptive event regardless of the adjusted CTCvalue. Some embodiments may implement other ways of dealing withmalevolent software, for example by aborting the transaction andprocessing a deferred disruptive event whenever a transaction's adjustedCTC value reaches zero prematurely.

If a deferral delay has expired at 460, then the processor may abort thetransaction at 475, process the deferred disruptive event at 485, andreturn to get the next instruction at 410. In some embodiments, the nextinstruction may indicate the start of a retry of the transaction abortedat 475. Although method 400 presents the handling of an expired deferraldelay as occurring only between execution of instructions at 460, insome embodiments, expired deferral delay handling may occur at any time,for example during execution of an instruction. The timing associatedwith handling expired deferral delays is implementation-specific.

If no deferral delay has expired at 460, or if a deferred disruptiveevent is not pending at 455, then the processor may determine whether adisruptive event has been identified at 465. Although method 400presents the handling of a disruptive event identification as occurringonly between execution of individual instructions at 465, in someembodiments, disruptive event identification handling may occur at anytime, for example during execution of an instruction. The timingassociated with handling disruptive event identifications isimplementation-specific. Furthermore, the types of disruptive eventsidentified at 465 may vary according to the specific embodiment. Forexample, a processor may be associated with two types of asynchronousinterrupts: architected (ISA) and micro-architected. Architectedinterrupts may include I/O interrupts and timer interrupts.Micro-architected interrupts may not be visible to software but mayinterrupt processor firmware. One, both, or neither of architected andmicro-architected interrupts may be identified at 465, depending on theimplementation.

If a disruptive event is not identified at 465, then the processor mayreturn to get the next instruction in the transaction at 425. If adisruptive event is identified at 465, then the processor may determinewhether predefined deferral criteria are satisfied for deferring thedisruptive event at 470. Deferral criteria may be relatively simple,such as a threshold for an adjusted CTC value. As previously discussed,the processor may periodically adjust the CTC value as the transactionprogresses, so that an internal adjusted CTC value represents anapproximation of how close the transaction is to completion at the timea disruptive event is identified. The method and timing of CTCadjustment may vary among embodiments. An instruction-based CTC valuemay be decremented after each instruction is completed. In someembodiments, for example where superscalar execution may provide for thecompletion of multiple instructions at the same cycle, aninstruction-based CTC value may be decreased by the number ofinstructions completed since the last CTC update. Furthermore, note thatthe term “execution” as used within this disclosure, when referring tothe execution of instructions, refers to the completion of instructions.In some embodiments, for example those supporting out-of-order orspeculative execution, more instructions may be technically executedthan are actually completed.

Transactions with an adjusted CTC value below the threshold may beconsidered close enough to completion to warrant deferral of thedisruptive event at 470, while transactions with an adjusted CTC valueabove the threshold may considered too far from completion to warrantdeferral of the disruptive event. Deferral criteria may be more complex,and may include different thresholds based on different types ofdisruptive events. For example, transactions with an adjusted CTC valuebelow an interrupt threshold may warrant deferral of a pendinginterrupt, while transactions with the same adjusted CTC value may notwarrant deferral of a conflicting memory access. Deferral criteria mayalso include a limit on the number of deferred disruptive events thatmay be pending simultaneously. For example, a first disruptive eventidentified while processing a transaction may be deferred, but a seconddisruptive event may result in a transactional abort.

If the deferral criteria are satisfied at 470, then the processor maydefer processing of the disruptive event at 480 and return to get thenext instruction in the transaction at 425. If the deferral criteria arenot satisfied at 470, then the processor may abort the transaction at475, process the disruptive event at 485, and return to get the nextinstruction at 410. In some embodiments, the next instruction mayindicate the start of a retry of the transaction aborted at 475.

An example method 500 of an embodiment for processing atime-to-completion transaction-begin instruction for a transaction isillustrated in FIG. 5. At start 505, a processor may process thetransaction-begin instruction to determine the initial CTC value at 540.For method 500, the initial CTC value is the approximate amount of timeneeded to complete the transaction. The processor may then start a timerat 545 for use in adjusting the CTC value as the transaction progresses.In some embodiments, the processor may note the value of an existingtimer rather than starting a timer, or may use some other method fortracking the progress of time.

The processor may then identify the next instruction in the transactionat 555. If the next instruction in the transaction identified at 555indicates that the transaction is complete at 558, then the processormay end processing of the transaction at 595 without identifying adisruptive event, and therefore without using the CTC value specified bythe transaction-begin instruction to perform hint-aided disruptive eventprocessing. If the transaction is not complete at 558, then theprocessor may execute or otherwise process the next instruction at 565and may adjust the CTC value at 570 to account for the time spentprocessing the transaction at 565. In some embodiments, the CTC valuemay not be adjusted to account for the progress of the transaction untila disruptive event is identified. The processor may then determinewhether a disruptive event has been identified at 585. If a disruptiveevent is not identified at 585, then the processor may identify the nextinstruction in the transaction at 555.

If a disruptive event is identified at 585, then the processor mayperform hint-aided processing of the disruptive event at 590 and may endprocessing of the transaction at 595. Hint-aided processing of thedisruptive event may be similar to method 400 described in FIG. 4, butis not necessarily a deferral of the disruptive event followed bycontinued processing of the transaction. Hint-aided processing of thedisruptive event may include adjusting the CTC value based on progressof the transaction, determining whether deferral criteria are satisfiedbased on the adjusted CTC value, and determining whether a deferraldelay has expired. The hint-aided protocol may or may not include adeferral of the disruptive event in conjunction with an attempt tocomplete the transaction, including executing one or more additionalinstructions within the transaction. The hint-aided protocol may or maynot include a transactional abort.

An example method 600 of an embodiment for processing aninstructions-to-completion transaction-begin instruction for atransaction is illustrated in FIG. 6. At start 605, a processor mayprocess the transaction-begin instruction to determine the initial CTCvalue at 640. For method 600, the initial CTC value is the approximatenumber of instructions remaining to complete the transaction. Theprocessor may then load an instruction counter with the CTC value at645. In some embodiments, the processor may note the value of anexisting instruction counter rather than loading a new counter, or mayuse some other method for tracking execution of instructions within thetransaction.

The processor may then identify the next instruction in the transactionat 655. If the next instruction in the transaction identified at 655indicates that the transaction is complete at 658, then the processormay end processing of the transaction at 695 without identifying adisruptive event, and therefore without using the CTC value specified bythe transaction-begin instruction to perform hint-aided disruptive eventprocessing. If the transaction is not complete at 658, then theprocessor may execute or otherwise process the next instruction at 665and may adjust the CTC value by decrementing the counter at 670 toaccount for the instruction processed at 665. In some embodiments, theCTC value may not be adjusted to account for the progress of thetransaction until a disruptive event is identified. The processor maythen determine whether a disruptive event has been identified at 685. Ifa disruptive event is not identified at 685, then the processor mayidentify the next instruction in the transaction at 655.

If a disruptive event is identified at 685, then the processor mayperform hint-aided processing of the disruptive event at 690 and may endprocessing of the transaction at 695. Hint-aided processing of thedisruptive event may be similar to method 400 described in FIG. 4, butis not necessarily a deferral of the disruptive event followed bycontinued processing of the transaction. Hint-aided processing of thedisruptive event may include adjusting the CTC value based on progressof the transaction, determining whether deferral criteria are satisfiedbased on the adjusted CTC value, and determining whether a deferraldelay has expired. The hint-aided protocol may or may not include adeferral of the disruptive event in conjunction with an attempt tocomplete the transaction, including executing one or more additionalinstructions within the transaction. The hint-aided protocol may or maynot include a transactional abort.

An example system 700 of an embodiment for managing transactional abortswith a transaction-begin instruction is illustrated in FIG. 7. Althoughsystem 700 is illustrated as a number of discrete units performing anumber of defined functions, any of these defined functions may beimplemented using more, fewer, or different discrete units than thoseshown in system 700.

Within processor 710, a transaction processing unit 715 may process atransaction-begin instruction 765, which loads a register 725 with a CTCvalue. CTC adjustment unit 720 receives information from transactionprocessing unit 715 about the progress of the transaction, and uses thisprogress information along with the value of CTC register 725 to adjustthe CTC value and update CTC register 725. For example, CTC adjustmentunit 720 may be a decrementing unit which decrements CTC register 725each time transaction processing unit 715 executes an instruction.

Disruptive event identification unit 735 notifies transaction processingunit 715, disruptive event processing unit 740, and deferraldetermination unit 730 whenever a disruptive event occurs duringprocessing of the transaction. Deferral determination unit 730 thendetermines whether the value in CTC register 725 satisfies deferralcriteria 755. If deferral criteria 755 is not satisfied, deferraldetermination unit 730 notifies transaction processing unit 715 to abortthe transaction and process the disruptive event. If deferral criteria755 is satisfied, deferral determination unit 730 notifies transactionprocessing unit 715 to defer processing the disruptive event and attemptto complete the transaction. Deferral determination unit 730 may alsonotify transaction processing unit 715 to abort the transaction andprocess a disruptive event if the disruptive event has been deferred forlonger than the deferral delay 760. Whenever transaction processing unit715 is notified to abort a transaction, it notifies transactional abortunit 750. Whenever transaction processing unit 715 is notified toprocess a disruptive event, it notifies disruptive event processing unit740.

Note that all instruction mnemonics used within this disclosure, such asTBEGIN, TEND, XBEGIN, and XACQUIRE, are selected for brevity ofreference and clarity of meaning and carry no special significance; anymnemonic could be selected to represent instructions or other entitiesin the disclosed computing environments.

Various embodiments of the present disclosure may be implemented in adata processing system suitable for storing and/or executing programcode that includes at least one processor coupled directly or indirectlyto memory elements through a system bus. The memory elements include,for instance, local memory employed during actual execution of theprogram code, bulk storage, and cache memory which provide temporarystorage of at least some program code in order to reduce the number oftimes code must be retrieved from bulk storage during execution.

Input/Output or I/O devices (including, but not limited to, keyboards,displays, pointing devices, DASD, tape, CDs, DVDs, thumb drives andother memory media, etc.) can be coupled to the system either directlyor through intervening I/O controllers. Network adapters may also becoupled to the system to enable the data processing system to becomecoupled to other data processing systems or remote printers or storagedevices through intervening private or public networks. Modems, cablemodems, and Ethernet cards are just a few of the available types ofnetwork adapters.

One or more of the capabilities of the present disclosure can beimplemented in software, firmware, hardware, or some combinationthereof. Further, one or more of the capabilities can be emulated. Forexample, an environment may include an emulator (e.g., software or otheremulation mechanisms), in which a particular architecture (including,for instance, instruction execution, architected functions, such asaddress translation, and architected registers) or a subset thereof isemulated (e.g., on a native computer system having a processor andmemory). In such an environment, one or more emulation functions of theemulator can implement one or more embodiments of the presentdisclosure, even though a computer executing the emulator may have adifferent architecture than the capabilities being emulated. As oneexample, in emulation mode, the specific instruction or operation beingemulated is decoded, and an appropriate emulation function is built toimplement the individual instruction or operation.

In FIG. 8, an example of an emulated host computer system 802 isprovided that emulates a host computer system 800 of a hostarchitecture. In the emulated host computer system 802, the hostprocessor (CPU) 810 is an emulated host processor (or virtual hostprocessor) and comprises an emulation processor 812 having a differentnative instruction set architecture than that of the processor 810 ofthe host computer 800. The emulated host computer system 802 has memory804 accessible to the emulation processor 812. In the exampleembodiment, the memory 804 is partitioned into a host computer memory806 portion and an emulation routines 808 portion. The host computermemory 806 is available to programs of the emulated host computer 802according to host computer architecture. The emulation processor 812executes native instructions of an architected instruction set of anarchitecture other than that of the emulated processor 810, the nativeinstructions obtained from emulation routines memory 808, and may accessa host instruction for execution from a program in host computer memory806 by employing one or more instructions obtained in a sequence &access/decode routine which may decode the host instruction(s) accessedto determine a native instruction execution routine for emulating thefunction of the host instruction accessed. Other facilities that aredefined for the host computer system 800 architecture may be emulated byarchitected facilities routines, including such facilities as generalpurpose registers, control registers, dynamic address translation andI/O subsystem support and processor cache, for example. The emulationroutines may also take advantage of functions available in the emulationprocessor 812 (such as general registers and dynamic translation ofvirtual addresses) to improve performance of the emulation routines.Special hardware and off-load engines may also be provided to assist theprocessor 812 in emulating the function of the host computer 800.

One or more aspects of the present disclosure can be included in anarticle of manufacture 900 (e.g., one or more computer program products)having, for instance, computer readable storage media 902 as shown inFIG. 9. The media has embodied therein, for instance, computer readableprogram code (instructions) 904 to provide and facilitate thecapabilities of the present disclosure. The article of manufacture 900can be included as a part of a computer system or as a separate product.

The present invention may be a system, a method, and/or a computerprogram product. The computer program product may include a computerreadable storage medium (or media) having computer readable programinstructions thereon for causing a processor to carry out aspects of thepresent invention.

The computer readable storage medium can be a tangible device that canretain and store instructions for use by an instruction executiondevice. The computer readable storage medium may be, for example, but isnot limited to, an electronic storage device, a magnetic storage device,an optical storage device, an electromagnetic storage device, asemiconductor storage device, or any suitable combination of theforegoing. A non-exhaustive list of more specific examples of thecomputer readable storage medium includes the following: a portablecomputer diskette, a hard disk, a random access memory (RAM), aread-only memory (ROM), an erasable programmable read-only memory (EPROMor Flash memory), a static random access memory (SRAM), a portablecompact disc read-only memory (CD-ROM), a digital versatile disk (DVD),a memory stick, a floppy disk, a mechanically encoded device such aspunch-cards or raised structures in a groove having instructionsrecorded thereon, and any suitable combination of the foregoing. Acomputer readable storage medium, as used herein, is not to be construedas being transitory signals per se, such as radio waves or other freelypropagating electromagnetic waves, electromagnetic waves propagatingthrough a waveguide or other transmission media (e.g., light pulsespassing through a fiber-optic cable), or electrical signals transmittedthrough a wire.

Computer readable program instructions described herein can bedownloaded to respective computing/processing devices from a computerreadable storage medium or to an external computer or external storagedevice via a network, for example, the Internet, a local area network, awide area network and/or a wireless network. The network may comprisecopper transmission cables, optical transmission fibers, wirelesstransmission, routers, firewalls, switches, gateway computers and/oredge servers. A network adapter card or network interface in eachcomputing/processing device receives computer readable programinstructions from the network and forwards the computer readable programinstructions for storage in a computer readable storage medium withinthe respective computing/processing device.

Computer readable program instructions for carrying out operations ofthe present invention may be assembler instructions,instruction-set-architecture (ISA) instructions, machine instructions,machine dependent instructions, microcode, firmware instructions,state-setting data, or either source code or object code written in anycombination of one or more programming languages, including an objectoriented programming language such as Java, Smalltalk, C++ or the like,and conventional procedural programming languages, such as the “C”programming language or similar programming languages. The computerreadable program instructions may execute entirely on the user'scomputer, partly on the user's computer, as a stand-alone softwarepackage, partly on the user's computer and partly on a remote computeror entirely on the remote computer or server. In the latter scenario,the remote computer may be connected to the user's computer through anytype of network, including a local area network (LAN) or a wide areanetwork (WAN), or the connection may be made to an external computer(for example, through the Internet using an Internet Service Provider).In some embodiments, electronic circuitry including, for example,programmable logic circuitry, field-programmable gate arrays (FPGA), orprogrammable logic arrays (PLA) may execute the computer readableprogram instructions by utilizing state information of the computerreadable program instructions to personalize the electronic circuitry,in order to perform aspects of the present invention.

Aspects of the present invention are described herein with reference toflowchart illustrations and/or block diagrams of methods, apparatus(systems), and computer program products according to embodiments of theinvention. It will be understood that each block of the flowchartillustrations and/or block diagrams, and combinations of blocks in theflowchart illustrations and/or block diagrams, can be implemented bycomputer readable program instructions.

These computer readable program instructions may be provided to aprocessor of a general purpose computer, special purpose computer, orother programmable data processing apparatus to produce a machine, suchthat the instructions, which execute via the processor of the computeror other programmable data processing apparatus, create means forimplementing the functions/acts specified in the flowchart and/or blockdiagram block or blocks. These computer readable program instructionsmay also be stored in a computer readable storage medium that can directa computer, a programmable data processing apparatus, and/or otherdevices to function in a particular manner, such that the computerreadable storage medium having instructions stored therein comprises anarticle of manufacture including instructions which implement aspects ofthe function/act specified in the flowchart and/or block diagram blockor blocks.

The computer readable program instructions may also be loaded onto acomputer, other programmable data processing apparatus, or other deviceto cause a series of operational steps to be performed on the computer,other programmable apparatus or other device to produce a computerimplemented process, such that the instructions which execute on thecomputer, other programmable apparatus, or other device implement thefunctions/acts specified in the flowchart and/or block diagram block orblocks.

The flowchart and block diagrams in the Figures illustrate thearchitecture, functionality, and operation of possible implementationsof systems, methods, and computer program products according to variousembodiments of the present invention. In this regard, each block in theflowchart or block diagrams may represent a module, segment, or portionof instructions, which comprises one or more executable instructions forimplementing the specified logical function(s). In some alternativeimplementations, the functions noted in the block may occur out of theorder noted in the figures. For example, two blocks shown in successionmay, in fact, be executed substantially concurrently, or the blocks maysometimes be executed in the reverse order, depending upon thefunctionality involved. It will also be noted that each block of theblock diagrams and/or flowchart illustration, and combinations of blocksin the block diagrams and/or flowchart illustration, can be implementedby special purpose hardware-based systems that perform the specifiedfunctions or acts or carry out combinations of special purpose hardwareand computer instructions.

Although one or more examples have been provided herein, these are onlyexamples. Many variations are possible without departing from the spiritof the present disclosure. For instance, processing environments otherthan the examples provided herein may include and/or benefit from one ormore aspects of the present disclosure. Further, the environment neednot be based on the z/Architecture®, but instead can be based on otherarchitectures offered by, for instance, IBM®, Intel®, Oracle®, as wellas others. Yet further, the environment can include multiple processors,be partitioned, and/or be coupled to other systems, as examples.

As used herein, the term “obtaining” includes, but is not limited to,fetching, receiving, having, providing, being provided, creating,developing, etc.

Although preferred embodiments have been depicted and described indetail herein, it will be apparent to those skilled in the relevant artthat various modifications, additions, substitutions and the like can bemade without departing from the spirit of the disclosure, and these are,therefore, considered to be within the scope of the disclosure, asdefined in the following claims.

What is claimed is:
 1. A method for processing transactions in atransaction execution (TX) computing environment with transactionalmemory, the method comprising: executing, by a processor, atransaction-begin instruction associated with a transaction, thetransaction-begin instruction specifying an initial value for atransaction-count-to-completion (CTC) value, the initial valueindicating a predicted duration of the transaction; adjusting, by theprocessor, the CTC value based on progress of the transaction, theadjusted CTC value predicting how far the transaction is fromcompletion; identifying, by the processor, a pending disruptive eventwhile processing the transaction, the disruptive event associated withinducing transactional aborts; based on the identified pendingdisruptive event, determining, by the processor, whether the adjustedCTC value satisfies predefined deferral criteria; and based ondetermining that the adjusted CTC value satisfies the predefineddeferral criteria, deferring processing of the identified pendingdisruptive event to avoid aborting the transaction.
 2. The method ofclaim 1, wherein an operand associated with the transaction-begininstruction specifies the initial value, and wherein the operand isselected from the group consisting of: an instructions-based operand,wherein the initial value indicates a predicted number of instructionsto execute in the transaction; a time-based operand, wherein the initialvalue indicates a predicted amount of time remaining to execute thetransaction; and a generic operand, wherein an indicator flag in a maskfield associated with the transaction-begin instruction specifieswhether the initial value indicates the predicted number of instructionsto execute in the transaction or the predicted amount of time remainingto execute the transaction.
 3. The method of claim 1, furthercomprising: based on determining that the adjusted CTC value does notsatisfy the predefined deferral criteria, aborting the transaction basedon the pending disruptive event.
 4. The method of claim 1, wherein thedisruptive event is selected from the group consisting of: a pendingasynchronous interrupt; and a conflicting memory access of memorymonitored by the transaction.
 5. The method of claim 1, wherein theinitial value is based on at least one of an immediate value associatedwith the transaction-begin instruction, a register identifier associatedwith the transaction-begin instruction, and a memory address associatedwith the transaction-begin instruction.